Speculation control for improving transaction success rate, and instruction therefor

ABSTRACT

Throttling instruction execution in a transaction operating in a processor configured to execute memory instructions out-of-order in a pipelined processor, wherein memory instructions are instructions for accessing operands in memory is provided. Included is executing, by the processor, instructions of a transaction comprising determining whether the transaction is in throttling mode and based on the transaction being in throttling mode, executing memory instructions in-program-order. Also included is based on the transaction not-being in throttling mode, executing memory instructions out-of-program order.

BACKGROUND

The present embodiment relates generally to transactional execution, andmore specifically to improving transaction success rate when speculationcontrol is enabled.

The number of central processing unit (CPU) cores on a chip and thenumber of CPU cores connected to a shared memory continues to growsignificantly to support growing workload capacity demand. Theincreasing number of CPUs cooperating to process the same workloads putsa significant burden on software scalability; for example, shared queuesor data-structures protected by traditional semaphores become hot spotsand lead to sub-linear n-way scaling curves. Traditionally this has beencountered by implementing finer-grained locking in software, and withlower latency/higher bandwidth interconnects in hardware. Implementingfine-grained locking to improve software scalability can be verycomplicated and error-prone, and at today's CPU frequencies, thelatencies of hardware interconnects are limited by the physicaldimension of the chips and systems, and by the speed of light.

Implementations of hardware Transactional Memory (TM) have beenintroduced, wherein a group of instructions, called a transaction,operate atomically and in isolation (sometimes called “serializability”)on a data structure in memory. The transaction executes optimisticallywithout obtaining a lock, but may need to abort and retry thetransaction execution if an operation, of the executing transaction, ona memory location conflicts with anther operation on the same memorylocation. Previously, software transactional memory implementations havebeen proposed to support software Transactional Memory (TM). However,hardware TM can provide improved performance aspects and ease of useover software TM.

U.S. Patent Application Publication No 2012/01599461 titled “ProgramOptimizing Apparatus, Program Optimizing Method, And Program OptimizingArticle Of Manufacture” filed 2012 Jun. 21 and incorporated by referenceherein teaches An apparatus having a transactional memory enablingexclusive control to execute a transaction. The apparatus includes: afirst code generating unit configured to interpret a program, andgenerate first code in which a begin instruction to begin a transactionand an end instruction to commit the transaction are inserted before andafter an instruction sequence including multiple instructions to executedesignated processing in the program; a second code generating unitconfigured to generate second code at a predetermined timing by usingthe multiple instructions according to the designated processing; and acode write unit configured to overwrite the instruction sequence of thefirst code with the second code or to write the second code to a part ofthe first code in the transaction.

U.S. Patent 2011/0246725 titled “System and Method for CommittingResults of a Software Transaction Using a Hardware Transaction” filed2011 Oct. 6 and incorporated by reference herein teaches The system andmethods described herein may exploit hardware transactional memory toimprove the performance of a software or hybrid transactional memoryimplementation, even when an entire user transaction cannot be executedwithin a hardware transaction. The user code of an atomic transactionmay be executed within a software transaction, which may collect readand write sets and/or other information about the atomic transaction. Asingle hardware transaction may be used to commit the atomic transactionby validating the transaction's read set and applying the effects of theuser code to memory, reducing the overhead associated with commitment ofsoftware transactions. Because the hardware transaction code iscarefully controlled, it may be less likely to fail to commit Variousremedial actions may be taken before retrying hardware transactionsfollowing some failures. If a transaction exceeds the constraints of thehardware, it may be committed by the software transactional memoryalone.

SUMMARY

Throttling instruction execution in a transaction operating in aprocessor configured to execute memory instructions out-of-order in apipelined processor, wherein memory instructions are instructions foraccessing operands in memory is provided. Included is executing, by theprocessor, instructions of a transaction comprising determining whetherthe transaction is in throttling mode and based on the transaction beingin throttling mode, executing memory instructions in-program-order. Alsoincluded is based on the transaction not-being in throttling mode,executing memory instructions out-of-program order.

BRIEF DESCRIPTION OF THE SEVERAL VIEWS OF THE DRAWINGS

These and other objects, features and advantages of the presentembodiment will become apparent from the following detailed descriptionof illustrative embodiments thereof, which is to be read in connectionwith the accompanying drawings. The various features of the drawings arenot to scale as the illustrations are for clarity in facilitating oneskilled in the art in understanding the embodiment in conjunction withthe detailed description. In the drawings:

FIGS. 1 and 2 depict an example multicore Transactional Memoryenvironment, in accordance with the present embodiment;

FIG. 3 depicts example components of an example CPU, in accordance withthe present embodiment;

FIG. 4 depicts a conventional processor with predictor update logic, inaccordance with the present embodiment;

FIG. 5 depicts an exemplary instruction tracking apparatus in anexemplary instruction sequencing unit (ISU), in accordance with thepresent embodiment;

FIG. 6 depicts an exemplary implementation of an issue queue andinstruction selection in instruction sequencing unit (ISU), inaccordance with the present embodiment;

FIG. 7 depicts an operational flowchart illustrating the operationscarried out by a program to improve memory performance when next tocomplete (NTC) speculation throttling is engaged, in accordance with thepresent embodiment;

FIG. 8 depicts a modified load and store unit (LSU) issue queue,speculating across memory instructions, in accordance with the presentembodiment;

FIGS. 9-11 depict an operational flowchart illustrating the logic for awindow control register, in accordance with the present embodiment;

FIG. 12 depicts a modified load and store unit (LSU) issue queue,speculating across memory instructions, in accordance with the presentembodiment;

FIG. 13 depicts a modified load and store unit (LSU) issue queuespeculating across memory instructions, in accordance with the presentembodiment;

FIG. 14 depicts a matrix scheduler in accordance with current methods;

FIG. 15 depicts a matrix scheduler executing instructions only after allprior instructions have executed, in accordance with the presentembodiment;

FIG. 16 depicts a matrix scheduler executing instructions only after allprior branch instructions have executed in order to verify correctnessof branch predictions, in accordance with the present embodiment;

FIG. 17 depicts an adaptivity flowchart illustrating the operationscarried out by a processor to determine whether the speculationthrottling mode is enabled in accordance with the present embodiment;

FIG. 18 depicts a flowchart illustrating the operations to enablethrottling speculation based on dynamic behavior in accordance with thepresent embodiment;

FIG. 19 depicts a flowchart illustrating the operations for a performprocessor assist (PPA) instruction in accordance with the presentembodiment;

FIG. 20 depicts an exemplary flowchart where throttling instructionexecution in a transaction operating in a processor configured toexecute memory instructions out-of-order in a pipelined processor,wherein memory instructions are instructions for accessing operands inmemory in accordance with the present embodiment; and

FIG. 21 depicts a block diagram of internal and external components ofcomputers and servers, in accordance with the present embodiment.

DETAILED DESCRIPTION

Historically, a computer system or processor had only a single processor(aka processing unit or central processing unit). The processor includedan instruction processing unit (IPU), a branch unit, a memory controlunit and the like. Such processors were capable of executing a singlethread of a program at a time. Operating systems were developed thatcould time-share a processor by dispatching a program to be executed onthe processor for a period of time, and then dispatching another programto be executed on the processor for another period of time. Astechnology evolved, memory subsystem caches were often added to theprocessor as well as complex dynamic address translation includingtranslation lookaside buffers (TLBs). The IPU itself was often referredto as a processor. As technology continued to evolve, an entireprocessor, could be packaged as a single semiconductor chip or die, sucha processor was referred to as a microprocessor. Then processors weredeveloped that incorporated multiple IPUs, such processors were oftenreferred to as multi-processors. Each such processor of amulti-processor computer system (processor) may include individual orshared caches, memory interfaces, system bus, address translationmechanism and the like. Virtual machine and instruction set architecture(ISA) emulators added a layer of software to a processor, that providedthe virtual machine with multiple “virtual processors” (aka processors)by time-slice usage of a single IPU in a single hardware processor. Astechnology further evolved, multi-threaded processors were developed,enabling a single hardware processor having a single multi-thread IPU toprovide a capability of simultaneously executing threads of differentprograms, thus each thread of a multi-threaded processor appeared to theoperating system as a processor. As technology further evolved, it waspossible to put multiple processors (each having an IPU) on a singlesemiconductor chip or die. These processors were referred to processorcores or just cores. Thus the terms such as processor, centralprocessing unit, processing unit, microprocessor, core, processor core,processor thread, and thread, for example, are often usedinterchangeably. Aspects of embodiments herein may be practiced by anyor all processors including those shown supra, without departing fromthe teachings herein. Wherein the term “thread” or “processor thread” isused herein, it is expected that particular advantage of the embodimentmay be had in a processor thread implementation.

Transaction Execution in Intel® Based Embodiments

In “Intel® Architecture Instruction Set Extensions ProgrammingReference” 319433-012A, February 2012, incorporated herein by referencein its entirety, Chapter 8 teaches, in part, that multithreadedapplications may take advantage of increasing numbers of CPU cores toachieve higher performance. However, the writing of multi-threadedapplications requires programmers to understand and take into accountdata sharing among the multiple threads. Access to shared data typicallyrequires synchronization mechanisms. These synchronization mechanismsare used to ensure that multiple threads update shared data byserializing operations that are applied to the shared data, oftenthrough the use of a critical section that is protected by a lock. Sinceserialization limits concurrency, programmers try to limit the overheaddue to synchronization.

Intel® Transactional Synchronization Extensions (Intel® TSX) allow aprocessor to dynamically determine whether threads need to be serializedthrough lock-protected critical sections, and to perform thatserialization only when required. This allows the processor to exposeand exploit concurrency that is hidden in an application because ofdynamically unnecessary synchronization.

With Intel TSX, programmer-specified code regions (also referred to as“transactional regions” or just “transactions”) are executedtransactionally. If the transactional execution completes successfully,then all memory operations performed within the transactional regionwill appear to have occurred instantaneously when viewed from otherprocessors. A processor makes the memory operations of the executedtransaction, performed within the transactional region, visible to otherprocessors only when a successful commit occurs, i.e., when thetransaction successfully completes execution. This process is oftenreferred to as an atomic commit.

Intel TSX provides two software interfaces to specify regions of codefor transactional execution. Hardware Lock Elision (HLE) is a legacycompatible instruction set extension (comprising the XACQUIRE andXRELEASE prefixes) to specify transactional regions. RestrictedTransactional Memory (RTM) is a new instruction set interface(comprising the XBEGIN, XEND, and XABORT instructions) for programmersto define transactional regions in a more flexible manner than thatpossible with HLE. HLE is for programmers who prefer the backwardcompatibility of the conventional mutual exclusion programming model andwould like to run HLE-enabled software on legacy hardware but would alsolike to take advantage of the new lock elision capabilities on hardwarewith HLE support. RTM is for programmers who prefer a flexible interfaceto the transactional execution hardware. In addition, Intel TSX alsoprovides an XTEST instruction. This instruction allows software to querywhether the logical processor is transactionally executing in atransactional region identified by either HLE or RTM.

Since a successful transactional execution ensures an atomic commit, theprocessor executes the code region optimistically without explicitsynchronization. If synchronization was unnecessary for that specificexecution, execution can commit without any cross-thread serialization.If the processor cannot commit atomically, then the optimistic executionfails. When this happens, the processor will roll back the execution, aprocess referred to as a transactional abort. On a transactional abort,the processor will discard all updates performed in the memory regionused by the transaction, restore architectural state to appear as if theoptimistic execution never occurred, and resume executionnon-transactionally.

A processor can perform a transactional abort for numerous reasons. Aprimary reason to abort a transaction is due to conflicting memoryaccesses between the transactionally executing logical processor andanother logical processor. Such conflicting memory accesses may preventa successful transactional execution. Memory addresses read from withina transactional region constitute the read-set of the transactionalregion and addresses written to within the transactional regionconstitute the write-set of the transactional region. Intel TSXmaintains the read- and write-sets at the granularity of a cache line. Aconflicting memory access occurs if another logical processor eitherreads a location that is part of the transactional region's write-set orwrites a location that is a part of either the read- or write-set of thetransactional region. A conflicting access typically means thatserialization is required for this code region. Since Intel TSX detectsdata conflicts at the granularity of a cache line, unrelated datalocations placed in the same cache line will be detected as conflictsthat result in transactional aborts. Transactional aborts may also occurdue to limited transactional resources. For example, the amount of dataaccessed in the region may exceed an implementation-specific capacity.Additionally, some instructions and system events may causetransactional aborts. Frequent transactional aborts result in wastedcycles and increased inefficiency.

Hardware Lock Elision

Hardware Lock Elision (HLE) provides a legacy compatible instruction setinterface for programmers to use transactional execution. HLE providestwo new instruction prefix hints: XACQUIRE and XRELEASE.

With HLE, a programmer adds the XACQUIRE prefix to the front of theinstruction that is used to acquire the lock that is protecting thecritical section. The processor treats the prefix as a hint to elide thewrite associated with the lock acquire operation. Even though the lockacquire has an associated write operation to the lock, the processordoes not add the address of the lock to the transactional region'swrite-set nor does it issue any write requests to the lock. Instead, theaddress of the lock is added to the read-set. The logical processorenters transactional execution. If the lock was available before theXACQUIRE prefixed instruction, then all other processors will continueto see the lock as available afterwards. Since the transactionallyexecuting logical processor neither added the address of the lock to itswrite-set nor performed externally visible write operations to the lock,other logical processors can read the lock without causing a dataconflict. This allows other logical processors to also enter andconcurrently execute the critical section protected by the lock. Theprocessor automatically detects any data conflicts that occur during thetransactional execution and will perform a transactional abort ifnecessary.

Even though the eliding processor did not perform any external writeoperations to the lock, the hardware ensures program order of operationson the lock. If the eliding processor itself reads the value of the lockin the critical section, it will appear as if the processor had acquiredthe lock, i.e. the read will return the non-elided value. This behaviorallows an HLE execution to be functionally equivalent to an executionwithout the HLE prefixes.

An XRELEASE prefix can be added in front of an instruction that is usedto release the lock protecting a critical section. Releasing the lockinvolves a write to the lock. If the instruction is to restore the valueof the lock to the value the lock had prior to the XACQUIRE prefixedlock acquire operation on the same lock, then the processor elides theexternal write request associated with the release of the lock and doesnot add the address of the lock to the write-set. The processor thenattempts to commit the transactional execution.

With HLE, if multiple threads execute critical sections protected by thesame lock but they do not perform any conflicting operations on eachother's data, then the threads can execute concurrently and withoutserialization. Even though the software uses lock acquisition operationson a common lock, the hardware recognizes this, elides the lock, andexecutes the critical sections on the two threads without requiring anycommunication through the lock—if such communication was dynamicallyunnecessary.

If the processor is unable to execute the region transactionally, thenthe processor will execute the region non-transactionally and withoutelision. HLE enabled software has the same forward progress guaranteesas the underlying non-HLE lock-based execution. For successful HLEexecution, the lock and the critical section code must follow certainguidelines. These guidelines only affect performance; and failure tofollow these guidelines will not result in a functional failure.Hardware without HLE support will ignore the XACQUIRE and XRELEASEprefix hints and will not perform any elision since these prefixescorrespond to the REPNE/REPE IA-32 prefixes which are ignored on theinstructions where XACQUIRE and XRELEASE are valid. Importantly, HLE iscompatible with the existing lock-based programming model. Improper useof hints will not cause functional bugs though it may expose latent bugsalready in the code.

Restricted Transactional Memory (RTM) provides a flexible softwareinterface for transactional execution. RTM provides three newinstructions—XBEGIN, XEND, and XABORT—for programmers to start, commit,and abort a transactional execution.

The programmer uses the XBEGIN instruction to specify the start of atransactional code region and the XEND instruction to specify the end ofthe transactional code region. If the RTM region could not besuccessfully executed transactionally, then the XBEGIN instruction takesan operand that provides a relative offset to the fallback instructionaddress.

A processor may abort RTM transactional execution for many reasons. Inmany instances, the hardware automatically detects transactional abortconditions and restarts execution from the fallback instruction addresswith the architectural state corresponding to that present at the startof the XBEGIN instruction and the EAX register updated to describe theabort status.

The XABORT instruction allows programmers to abort the execution of anRTM region explicitly. The XABORT instruction takes an 8-bit immediateargument that is loaded into the EAX register and will thus be availableto software following an RTM abort. RTM instructions do not have anydata memory location associated with them. While the hardware providesno guarantees as to whether an RTM region will ever successfully committransactionally, most transactions that follow the recommendedguidelines are expected to successfully commit transactionally. However,programmers must always provide an alternative code sequence in thefallback path to guarantee forward progress. This may be as simple asacquiring a lock and executing the specified code regionnon-transactionally. Further, a transaction that always aborts on agiven implementation may complete transactionally on a futureimplementation. Therefore, programmers must ensure the code paths forthe transactional region and the alternative code sequence arefunctionally tested.

Detection of HLE Support

A processor supports HLE execution if CPUID.07H.EBX.HLE [bit 4]=1.However, an application can use the HLE prefixes (XACQUIRE and XRELEASE)without checking whether the processor supports HLE. Processors withoutHLE support ignore these prefixes and will execute the code withoutentering transactional execution.

Detection of RTM Support

A processor supports RTM execution if CPUID.07H.EBX.RTM [bit 11]=1. Anapplication must check if the processor supports RTM before it uses theRTM instructions (XBEGIN, XEND, XABORT). These instructions willgenerate a #UD exception when used on a processor that does not supportRTM.

Detection of XTEST Instruction

A processor supports the XTEST instruction if it supports either HLE orRTM. An application must check either of these feature flags beforeusing the XTEST instruction. This instruction will generate a #UDexception when used on a processor that does not support either HLE orRTM.

Querying Transactional Execution Status

The XTEST instruction can be used to determine the transactional statusof a transactional region specified by HLE or RTM. Note, while the HLEprefixes are ignored on processors that do not support HLE, the XTESTinstruction will generate a #UD exception when used on processors thatdo not support either HLE or RTM.

Requirements for HLE Locks

For HLE execution to successfully commit transactionally, the lock mustsatisfy certain properties and access to the lock must follow certainguidelines.

An XRELEASE prefixed instruction must restore the value of the elidedlock to the value it had before the lock acquisition. This allowshardware to safely elide locks by not adding them to the write-set. Thedata size and data address of the lock release (XRELEASE prefixed)instruction must match that of the lock acquire (XACQUIRE prefixed) andthe lock must not cross a cache line boundary.

Software should not write to the elided lock inside a transactional HLEregion with any instruction other than an XRELEASE prefixed instruction,otherwise such a write may cause a transactional abort. In addition,recursive locks (where a thread acquires the same lock multiple timeswithout first releasing the lock) may also cause a transactional abort.Note that software can observe the result of the elided lock acquireinside the critical section. Such a read operation will return the valueof the write to the lock.

The processor automatically detects violations to these guidelines, andsafely transitions to a non-transactional execution without elision.Since Intel TSX detects conflicts at the granularity of a cache line,writes to data collocated on the same cache line as the elided lock maybe detected as data conflicts by other logical processors eliding thesame lock.

Transactional Nesting

Both HLE and RTM support nested transactional regions. However, atransactional abort restores state to the operation that startedtransactional execution: either the outermost XACQUIRE prefixed HLEeligible instruction or the outermost XBEGIN instruction. The processortreats all nested transactions as one transaction.

HLE Nesting and Elision

Programmers can nest HLE regions up to an implementation specific depthof MAX_HLE_NEST_COUNT. Each logical processor tracks the nesting countinternally but this count is not available to software. An XACQUIREprefixed HLE-eligible instruction increments the nesting count, and anXRELEASE prefixed HLE-eligible instruction decrements it. The logicalprocessor enters transactional execution when the nesting count goesfrom zero to one. The logical processor attempts to commit only when thenesting count becomes zero. A transactional abort may occur if thenesting count exceeds MAX_HLE_NEST_COUNT.

In addition to supporting nested HLE regions, the processor can alsoelide multiple nested locks. The processor tracks a lock for elisionbeginning with the XACQUIRE prefixed HLE eligible instruction for thatlock and ending with the XRELEASE prefixed HLE eligible instruction forthat same lock. The processor can, at any one time, track up to aMAX_HLE_ELIDED_LOCKS number of locks. For example, if the implementationsupports a MAX_HLE_ELIDED_LOCKS value of two and if the programmer neststhree HLE identified critical sections (by performing XACQUIRE prefixedHLE eligible instructions on three distinct locks without performing anintervening XRELEASE prefixed HLE eligible instruction on any one of thelocks), then the first two locks will be elided, but the third won't beelided (but will be added to the transaction's writeset). However, theexecution will still continue transactionally. Once an XRELEASE for oneof the two elided locks is encountered, a subsequent lock acquiredthrough the XACQUIRE prefixed HLE eligible instruction will be elided.

The processor attempts to commit the HLE execution when all elidedXACQUIRE and XRELEASE pairs have been matched, the nesting count goes tozero, and the locks have satisfied requirements. If execution cannotcommit atomically, then execution transitions to a non-transactionalexecution without elision as if the first instruction did not have anXACQUIRE prefix.

RTM Nesting

Programmers can nest RTM regions up to an implementation specificMAX_RTM_NEST_COUNT. The logical processor tracks the nesting countinternally but this count is not available to software. An XBEGINinstruction increments the nesting count, and an XEND instructiondecrements the nesting count. The logical processor attempts to commitonly if the nesting count becomes zero. A transactional abort occurs ifthe nesting count exceeds MAX_RTM_NEST_COUNT.

Nesting HLE and RTM

HLE and RTM provide two alternative software interfaces to a commontransactional execution capability. Transactional processing behavior isimplementation specific when HLE and RTM are nested together, e.g., HLEis inside RTM or RTM is inside HLE. However, in all cases, theimplementation will maintain HLE and RTM semantics. An implementationmay choose to ignore HLE hints when used inside RTM regions, and maycause a transactional abort when RTM instructions are used inside HLEregions. In the latter case, the transition from transactional tonon-transactional execution occurs seamlessly since the processor willre-execute the HLE region without actually doing elision, and thenexecute the RTM instructions.

Abort Status Definition

RTM uses the EAX register to communicate abort status to software.Following an RTM abort the EAX register has the following definition.

TABLE 1 RTM Abort Status Definition EAX Register Bit Position Meaning 0Set if abort caused by XABORT instruction 1 If set, the transaction maysucceed on retry, this bit is always clear if bit 0 is set 2 Set ifanother logical processor conflicted with a memory address that was partof the transaction that aborted 3 Set if an internal buffer overflowed 4Set if a debug breakpoint was hit 5 Set if an abort occurred duringexecution of a nested transaction 23:6 Reserved 31-24 XABORT argument(only valid if bit 0 set, otherwise reserved)

The EAX abort status for RTM only provides causes for aborts. It doesnot by itself encode whether an abort or commit occurred for the RTMregion. The value of EAX can be 0 following an RTM abort. For example, aCPUID instruction when used inside an RTM region causes a transactionalabort and may not satisfy the requirements for setting any of the EAXbits. This may result in an EAX value of 0.

RTM Memory Ordering

A successful RTM commit causes all memory operations in the RTM regionto appear to execute atomically. A successfully committed RTM regionconsisting of an XBEGIN followed by an XEND, even with no memoryoperations in the RTM region, has the same ordering semantics as a LOCKprefixed instruction.

The XBEGIN instruction does not have fencing semantics. However, if anRTM execution aborts, then all memory updates from within the RTM regionare discarded and are not made visible to any other logical processor.

RTM-Enabled Debugger Support

By default, any debug exception inside an RTM region will cause atransactional abort and will redirect control flow to the fallbackinstruction address with architectural state recovered and bit 4 in EAXset. However, to allow software debuggers to intercept execution ondebug exceptions, the RTM architecture provides additional capability.

If bit 11 of DR7 and bit 15 of the IA32_DEBUGCTL_MSR are both 1, any RTMabort due to a debug exception (#DB) or breakpoint exception (#BP)causes execution to roll back and restart from the XBEGIN instructioninstead of the fallback address. In this scenario, the EAX register willalso be restored back to the point of the XBEGIN instruction.

Programming Considerations

Typical programmer-identified regions are expected to transactionallyexecute and commit successfully. However, Intel TSX does not provide anysuch guarantee. A transactional execution may abort for many reasons. Totake full advantage of the transactional capabilities, programmersshould follow certain guidelines to increase the probability of theirtransactional execution committing successfully.

This section discusses various events that may cause transactionalaborts. The architecture ensures that updates performed within atransaction that subsequently aborts execution will never becomevisible. Only committed transactional executions initiate an update tothe architectural state. Transactional aborts never cause functionalfailures and only affect performance.

Instruction Based Considerations

Programmers can use any instruction safely inside a transaction (HLE orRTM) and can use transactions at any privilege level. However, someinstructions will always abort the transactional execution and causeexecution to seamlessly and safely transition to a non-transactionalpath.

Intel TSX allows for most common instructions to be used insidetransactions without causing aborts. The following operations inside atransaction do not typically cause an abort:

-   -   Operations on the instruction pointer register, general purpose        registers (GPRs) and the status flags (CF, OF, SF, PF, AF, and        ZF); and    -   Operations on XMM and YMM registers and the MXCSR register.

However, programmers must be careful when intermixing SSE and AVXoperations inside a transactional region. Intermixing SSE instructionsaccessing XMM registers and AVX instructions accessing YMM registers maycause transactions to abort. Programmers may use REP/REPNE prefixedstring operations inside transactions. However, long strings may causeaborts. Further, the use of CLD and STD instructions may cause aborts ifthey change the value of the DF flag. However, if DF is 1, the STDinstruction will not cause an abort. Similarly, if DF is 0, then the CLDinstruction will not cause an abort.

Instructions not enumerated here as causing abort when used inside atransaction will typically not cause a transaction to abort (examplesinclude but are not limited to MFLNCE, LFLNCE, SFENCE, RDTSC, RDTSCP,etc.).

The following instructions will abort transactional execution on anyimplementation:

-   -   XABORT    -   CPUID    -   PAUSE

In addition, in some implementations, the following instructions mayalways cause transactional aborts. These instructions are not expectedto be commonly used inside typical transactional regions. However,programmers must not rely on these instructions to force a transactionalabort, since whether they cause transactional aborts is implementationdependent.

-   -   Operations on X87 and MMX architecture state. This includes all        MMX and X87 instructions, including the FXRSTOR and FXSAVE        instructions.    -   Update to non-status portion of EFLAGS: CLI, STI, POPFD, POPFQ,        CLTS.    -   Instructions that update segment registers, debug registers        and/or control registers: MOV to DS/ES/FS/GS/SS, POP        DS/ES/FS/GS/SS, LDS, LES, LFS, LGS, LSS, SWAPGS, WRFSBASE,        WRGSBASE, LGDT, SGDT, LIDT, SIDT, LLDT, SLDT, LTR, STR, Far        CALL, Far JMP, Far RET, IRET, MOV to DRx, MOV to        CRO/CR2/CR3/CR4/CR8 and LMSW.    -   Ring transitions: SYSENTER, SYSCALL, SYSEXIT, and SYSRET.    -   TLB and Cacheability control: CLFLUSH, INVD, WBINVD, INVLPG,        INVPCID, and memory instructions with a non-temporal hint        (MOVNTDQA, MOVNTDQ, MOVNTI, MOVNTPD, MOVNTPS, and MOVNTQ).    -   Processor state save: XSAVE, XSAVEOPT, and XRSTOR.    -   Interrupts: INTn, INTO.    -   IO: IN, INS, REP INS, OUT, OUTS, REP OUTS and their variants.    -   VMX: VMPTRLD, VMPTRST, VMCLEAR, VMREAD, VMWRITE, VMCALL,        VMLAUNCH, VMRESUME, VMXOFF, VMXON, INVEPT, and INVVPID.    -   SMX: GETSEC.    -   UD2, RSM, RDMSR, WRMSR, HLT, MONITOR, MWAIT, XSETBV, VZEROUPPER,        MASKMOVQ, and V/MASKMOVDQU.        Runtime Considerations

In addition to the instruction-based considerations, runtime events maycause transactional execution to abort. These may be due to data accesspatterns or micro-architectural implementation features. The followinglist is not a comprehensive discussion of all abort causes.

Any fault or trap in a transaction that must be exposed to software willbe suppressed. Transactional execution will abort and execution willtransition to a non-transactional execution, as if the fault or trap hadnever occurred. If an exception is not masked, then that un-maskedexception will result in a transactional abort and the state will appearas if the exception had never occurred.

Synchronous exception events (#DE, #OF, #NP, #SS, #GP, #BR, #UD, #AC,#XF, #PF, #NM, #TS, #MF, #DB, #BP/INT3) that occur during transactionalexecution may cause an execution not to commit transactionally, andrequire a non-transactional execution. These events are suppressed as ifthey had never occurred. With HLE, since the non-transactional code pathis identical to the transactional code path, these events will typicallyre-appear when the instruction that caused the exception is re-executednon-transactionally, causing the associated synchronous events to bedelivered appropriately in the non-transactional execution. Asynchronousevents (NMI, SMI, INTR, IPI, PMI, etc.) occurring during transactionalexecution may cause the transactional execution to abort and transitionto a non-transactional execution. The asynchronous events will be pendedand handled after the transactional abort is processed.

Transactions only support write-back cacheable memory type operations. Atransaction may always abort if the transaction includes operations onany other memory type. This includes instruction fetches to UC memorytype.

Memory accesses within a transactional region may require the processorto set the Accessed and Dirty flags of the referenced page table entry.The behavior of how the processor handles this is implementationspecific. Some implementations may allow the updates to these flags tobecome externally visible even if the transactional region subsequentlyaborts. Some Intel TSX implementations may choose to abort thetransactional execution if these flags need to be updated. Further, aprocessor's page-table walk may generate accesses to its owntransactionally written but uncommitted state. Some Intel TSXimplementations may choose to abort the execution of a transactionalregion in such situations. Regardless, the architecture ensures that, ifthe transactional region aborts, then the transactionally written statewill not be made architecturally visible through the behavior ofstructures such as TLBs.

Executing self-modifying code transactionally may also causetransactional aborts. Programmers must continue to follow the Intelrecommended guidelines for writing self-modifying and cross-modifyingcode even when employing HLE and RTM. While an implementation of RTM andHLE will typically provide sufficient resources for executing commontransactional regions, implementation constraints and excessive sizesfor transactional regions may cause a transactional execution to abortand transition to a non-transactional execution. The architectureprovides no guarantee of the amount of resources available to dotransactional execution and does not guarantee that a transactionalexecution will ever succeed.

Conflicting requests to a cache line accessed within a transactionalregion may prevent the transaction from executing successfully. Forexample, if logical processor P0 reads line A in a transactional regionand another logical processor P1 writes line A (either inside or outsidea transactional region) then logical processor P0 may abort if logicalprocessor P1's write interferes with processor P0's ability to executetransactionally.

Similarly, if P0 writes line A in a transactional region and P1 reads orwrites line A (either inside or outside a transactional region), then P0may abort if P1's access to line A interferes with P0's ability toexecute transactionally. In addition, other coherence traffic may attimes appear as conflicting requests and may cause aborts. While thesefalse conflicts may happen, they are expected to be uncommon. Theconflict resolution policy to determine whether P0 or P1 aborts in theabove scenarios is implementation specific.

Generic Transaction Execution embodiments:

According to “ARCHITECTURES FOR TRANSACTIONAL MEMORY”, a dissertationsubmitted to the Department of Computer Science and the Committee onGraduate Studies of Stanford University in partial fulfillment of therequirements for the Degree of Doctor of Philosophy, by Austen McDonald,June 2009, incorporated by reference herein in its entirety,fundamentally, there are three mechanisms needed to implement an atomicand isolated transactional region: versioning, conflict detection, andcontention management.

To make a transactional code region appear atomic, all the modificationsperformed by that transactional code region must be stored and keptisolated from other transactions until commit time. The system does thisby implementing a versioning policy. Two versioning paradigms exist:eager and lazy. An eager versioning system stores newly generatedtransactional values in place and stores previous memory values on theside, in what is called an undo-log. A lazy versioning system stores newvalues temporarily in what is called a write buffer, copying them tomemory only on commit. In either system, the cache is used to optimizestorage of new versions.

To ensure serializability between transactions, conflicts must bedetected and resolved. The two systems, i.e., the eager and lazyversioning systems, detect conflicts by implementing a conflictdetection policy, either optimistic or pessimistic. An optimistic systemexecutes transactions in parallel, checking for conflicts only when atransaction commits A pessimistic system checks for conflicts at eachload and store. Similar to versioning, conflict detection also uses thecache, marking each line as either part of the read-set, part of thewrite-set, or both. The two systems resolve conflicts by implementing acontention management policy. Many contention management policies exist,some are more appropriate for optimistic conflict detection and some aremore appropriate for pessimistic. Described below are some examplepolicies.

Since each transactional memory (TM) system needs both versioningdetection and conflict detection, these options give rise to fourdistinct TM designs: Eager-Pessimistic (EP), Eager-Optimistic (EO),Lazy-Pessimistic (LP), and Lazy-Optimistic (LO). Table 2 brieflydescribes all four distinct TM designs.

FIGS. 1 and 2 depict an example of a multicore TM environment. FIG. 1shows many TM-enabled CPUs (CPU1 114 a, CPU2 114 b, etc.) on one die100, connected with an interconnect 122, under management of aninterconnect control 120 a, 120 b. Each CPU 114 a, 114 b (also known asa Processor) may have a split cache consisting of an Instruction Cache116 a, 116 b for caching instructions from memory to be executed and aData Cache 118 a, 118 b with TM support for caching data (operands) ofmemory locations to be operated on by CPU 114 a, 114 b (in FIG. 1, eachCPU 114 a, 114 b and its associated caches are referenced as 112 a, 112b). In an implementation, caches of multiple dies 100 are interconnectedto support cache coherency between the caches of the multiple dies 100.In an implementation, a single cache, rather than the split cache isemployed holding both instructions and data. In implementations, the CPUcaches are one level of caching in a hierarchical cache structure. Forexample each die 100 may employ a shared cache 124 to be shared amongstall the CPUs on the die 100. In another implementation, each die mayhave access to a shared cache 124, shared amongst all the processors ofall the dies 100.

FIG. 2 shows the details of an example transactional CPU environment112, having a CPU 114, including additions to support TM. Thetransactional CPU (processor) 114 may include hardware for supportingRegister Checkpoints 126 and special TM Registers 128. The transactionalCPU cache may have the MESI bits 130, Tags 140 and Data 142 of aconventional cache but also, for example, R bits 132 showing a line hasbeen read by the CPU 114 while executing a transaction and W bits 138showing a line has been written-to by the CPU 114 while executing atransaction.

A key detail for programmers in any TM system is how non-transactionalaccesses interact with transactions. By design, transactional accessesare screened from each other using the mechanisms above. However, theinteraction between a regular, non-transactional load with a transactioncontaining a new value for that address must still be considered. Inaddition, the interaction between a non-transactional store with atransaction that has read that address must also be explored. These areissues of the database concept isolation.

A TM system is said to implement strong isolation, sometimes calledstrong atomicity, when every non-transactional load and store acts likean atomic transaction. Therefore, non-transactional loads cannot seeuncommitted data and non-transactional stores cause atomicity violationsin any transactions that have read that address. A system where this isnot the case is said to implement weak isolation, sometimes called weakatomicity.

Strong isolation is often more desirable than weak isolation due to therelative ease of conceptualization and implementation of strongisolation. Additionally, if a programmer has forgotten to surround someshared memory references with transactions, causing bugs, then withstrong isolation, the programmer will often detect that oversight usinga simple debug interface because the programmer will see anon-transactional region causing atomicity violations. Also, programswritten in one model may work differently on another model.

Further, strong isolation is often easier to support in hardware TM thanweak isolation. With strong isolation, since the coherence protocolalready manages load and store communication between processors,transactions can detect non-transactional loads and stores and actappropriately. To implement strong isolation in software TransactionalMemory (TM), non-transactional code must be modified to include read-and write-barriers; potentially crippling performance. Although greateffort has been expended to remove many un-needed barriers, suchtechniques are often complex and performance is typically far lower thanthat of HTMs.

TABLE 2 Transactional Memory Design Space VERSIONING Lazy Eager CONFLICTOptimistic Storing updates in a write buffer; Not practical: waiting toupdate memory DETECTION detecting conflicts at commit time. until committime but detecting conflicts at access time guarantees wasted work andprovides no advantage Pessimistic Storing updates in a write buffer;Updating memory, keeping old values in detecting conflicts at accesstime. undo log; detecting conflicts at access time.

Table 2 illustrates the fundamental design space of transactional memory(versioning and conflict detection).

Eager-Pessimistic (EP)

This first TM design described below is known as Eager-Pessimistic. AnEP system stores its write-set “in place” (hence the name “eager”) and,to support rollback, stores the old values of overwritten lines in an“undo log”. Processors use the W 138 and R 132 cache bits to track readand write-sets and detect conflicts when receiving snooped loadrequests. Perhaps the most notable examples of EP systems in knownliterature are LogTM and UTM.

Beginning a transaction in an EP system is much like beginning atransaction in other systems: tm_begin( ) takes a register checkpoint,and initializes any status registers. An EP system also requiresinitializing the undo log, the details of which are dependent on the logformat, but often involve initializing a log base pointer to a region ofpre-allocated, thread-private memory, and clearing a log boundsregister.

Versioning: In EP, due to the way eager versioning is designed tofunction, the MESI 130 state transitions (cache line indicatorscorresponding to Modified, Exclusive, Shared, and Invalid code states)are left mostly unchanged. Outside of a transaction, the MESI 130 statetransitions are left completely unchanged. When reading a line inside atransaction, the standard coherence transitions apply (S (Shared)→S, I(Invalid)→S, or I→E (Exclusive)), issuing a load miss as needed, but theR 132 bit is also set. Likewise, writing a line applies the standardtransitions (S→M, E→I, I→M), issuing a miss as needed, but also sets theW 138 (Written) bit. The first time a line is written, the old versionof the entire line is loaded then written to the undo log to preserve itin case the current transaction aborts. The newly written data is thenstored “in-place,” over the old data.

Conflict Detection: Pessimistic conflict detection uses coherencemessages exchanged on misses, or upgrades, to look for conflicts betweentransactions. When a read miss occurs within a transaction, otherprocessors receive a load request; but they ignore the request if theydo not have the needed line. If the other processors have the neededline non-speculatively or have the line R 132 (Read), they downgradethat line to S, and in certain cases issue a cache-to-cache transfer ifthey have the line in MESI's 130 M or E state. However, if the cache hasthe line W 138, then a conflict is detected between the two transactionsand additional action(s) must be taken.

Similarly, when a transaction seeks to upgrade a line from shared tomodified (on a first write), the transaction issues an exclusive loadrequest, which is also used to detect conflicts. If a receiving cachehas the line non-speculatively, then the line is invalidated, and incertain cases a cache-to-cache transfer (M or E states) is issued. But,if the line is R 132 or W 138, a conflict is detected.

Validation: Because conflict detection is performed on every load, atransaction always has exclusive access to its own write-set. Therefore,validation does not require any additional work.

Commit: Since eager versioning stores the new version of data items inplace, the commit process simply clears the W 138 and R 132 bits anddiscards the undo log.

Abort: When a transaction rolls back, the original version of each cacheline in the undo log must be restored, a process called “unrolling” or“applying” the log. This is done during tm_discard( ) and must be atomicwith regard to other transactions. Specifically, the write-set muststill be used to detect conflicts: this transaction has the only correctversion of lines in its undo log, and requesting transactions must waitfor the correct version to be restored from that log. Such a log can beapplied using a hardware state machine or software abort handler.

Eager-Pessimistic has the characteristics of: Commit is simple and sinceit is in-place, very fast. Similarly, validation is a no-op. Pessimisticconflict detection detects conflicts early, thereby reducing the numberof “doomed” transactions. For example, if two transactions are involvedin a Write-After-Read dependency, then that dependency is detectedimmediately in pessimistic conflict detection. However, in optimisticconflict detection such conflicts are not detected until the writercommits.

Eager-Pessimistic also has the characteristics of: As described above,the first time a cache line is written, the old value must be written tothe log, incurring extra cache accesses. Aborts are expensive as theyrequire undoing the log. For each cache line in the log, a load must beissued, perhaps going as far as main memory before continuing to thenext line. Pessimistic conflict detection also prevents certainserializable schedules from existing.

Additionally, because conflicts are handled as they occur, there is apotential for livelock and careful contention management mechanisms mustbe employed to guarantee forward progress.

Lazy-Optimistic (LO)

Another popular TM design is Lazy-Optimistic (LO), which stores itswrite-set in a “write buffer” or “redo log” and detects conflicts atcommit time (still using the R 132 and W 138 bits).

Versioning: Just as in the EP system, the MESI protocol of the LO designis enforced outside of the transactions. Once inside a transaction,reading a line incurs the standard MESI transitions but also sets the R132 bit. Likewise, writing a line sets the W 138 bit of the line, buthandling the MESI transitions of the LO design is different from that ofthe EP design. First, with lazy versioning, the new versions of writtendata are stored in the cache hierarchy until commit while othertransactions have access to old versions available in memory or othercaches. To make available the old versions, dirty lines (M lines) mustbe evicted when first written by a transaction. Second, no upgrademisses are needed because of the optimistic conflict detection feature:if a transaction has a line in the S state, it can simply write to itand upgrade that line to an M state without communicating the changeswith other transactions because conflict detection is done at committime.

Conflict Detection and Validation: To validate a transaction and detectconflicts, LO communicates the addresses of speculatively modified linesto other transactions only when it is preparing to commit. Onvalidation, the processor sends one, potentially large, network packetcontaining all the addresses in the write-set. Data is not sent, butleft in the cache of the committer and marked dirty (M). To build thispacket without searching the cache for lines marked W, a simple bitvector is used, called a “store buffer,” with one bit per cache line totrack these speculatively modified lines. Other transactions use thisaddress packet to detect conflicts: if an address is found in the cacheand the R 132 and/or W 138 bits are set, then a conflict is initiated.If the line is found but neither R 132 nor W 138 is set, then the lineis simply invalidated, which is similar to processing an exclusive load.

To support transaction atomicity, these address packets must be handledatomically, i.e., no two address packets may exist at once with the sameaddresses. In an LO system, this can be achieved by simply acquiring aglobal commit token before sending the address packet. However, atwo-phase commit scheme could be employed by first sending out theaddress packet, collecting responses, enforcing an ordering protocol(perhaps oldest transaction first), and committing once all responsesare satisfactory.

Commit Once validation has occurred, commit needs no special treatment:simply clear W 138 and R 132 bits and the store buffer. Thetransaction's writes are already marked dirty in the cache and othercaches' copies of these lines have been invalidated via the addresspacket. Other processors can then access the committed data through theregular coherence protocol.

Abort: Rollback is equally easy: because the write-set is containedwithin the local caches, these lines can be invalidated, then clear W138 and R 132 bits and the store buffer. The store buffer allows W linesto be found to invalidate without the need to search the cache.

Lazy-Optimistic has the characteristics of: Aborts are very fast,requiring no additional loads or stores and making only local changes.More serializable schedules can exist than found in EP, which allows anLO system to more aggressively speculate that transactions areindependent, which can yield higher performance Finally, the latedetection of conflicts can increase the likelihood of forward progress.

Lazy-Optimistic also has the characteristics of: Validation takes globalcommunication time proportional to size of write set. Doomedtransactions can waste work since conflicts are detected only at committime.

Lazy-Pessimistic (LP)

Lazy-Pessimistic (LP) represents a third TM design option, sittingsomewhere between EP and LO: storing newly written lines in a writebuffer but detecting conflicts on a per access basis.

Versioning: Versioning is similar but not identical to that of LO:reading a line sets its R bit 132, writing a line sets its W bit 138,and a store buffer is used to track W lines in the cache. Also, dirty(M) lines must be evicted when first written by a transaction, just asin LO. However, since conflict detection is pessimistic, load exclusivesmust be performed when upgrading a transactional line from I, S→M, whichis unlike LO.

Conflict Detection: LP's conflict detection operates the same as EP's:using coherence messages to look for conflicts between transactions.

Validation: Like in EP, pessimistic conflict detection ensures that atany point, a running transaction has no conflicts with any other runningtransaction, so validation is a no-op.

Commit: Commit needs no special treatment: simply clear W 138 and R 132bits and the store buffer, like in LO.

Abort: Rollback is also like that of LO: simply invalidate the write-setusing the store buffer and clear the W and R bits and the store buffer.

Eager-Optimistic (EO)

The LP has the characteristics of: Like LO, aborts are very fast. LikeEP, the use of pessimistic conflict detection reduces the number of“doomed” transactions. Like EP, some serializable schedules are notallowed and conflict detection must be performed on each cache miss.

The final combination of versioning and conflict detection isEager-Optimistic (EO). EO may be a less than optimal choice for HTMsystems: since new transactional versions are written in-place, othertransactions have no choice but to notice conflicts as they occur (i.e.,as cache misses occur). But since EO waits until commit time to detectconflicts, those transactions become “zombies,” continuing to execute,wasting resources, yet are “doomed” to abort.

EO has proven to be useful in STMs and is implemented by Bartok-STM andMcRT. A lazy versioning STM needs to check its write buffer on each readto ensure that it is reading the most recent value. Since the writebuffer is not a hardware structure, this is expensive, hence thepreference for write-in-place eager versioning. Additionally, sincechecking for conflicts is also expensive in an STM, optimistic conflictdetection offers the advantage of performing this operation in bulk.

Contention Management

How a transaction rolls back once the system has decided to abort thattransaction has been described above, but, since a conflict involves twotransactions, the topics of which transaction should abort, how thatabort should be initiated, and when should the aborted transaction beretried need to be explored. These are topics that are addressed byContention Management (CM), a key component of transactional memory.Described below are policies regarding how the systems initiate abortsand the various established methods of managing which transactionsshould abort in a conflict.

Contention Management Policies

A Contention Management (CM) Policy is a mechanism that determines whichtransaction involved in a conflict should abort and when the abortedtransaction should be retried. For example, it is often the case thatretrying an aborted transaction immediately does not lead to the bestperformance. Conversely, employing a back-off mechanism, which delaysthe retrying of an aborted transaction, can yield better performance.STMs first grappled with finding the best contention management policiesand many of the policies outlined below were originally developed forSTMs.

CM Policies draw on a number of measures to make decisions, includingages of the transactions, size of read- and write-sets, the number ofprevious aborts, etc. The combinations of measures to make suchdecisions are endless, but certain combinations are described below,roughly in order of increasing complexity.

To establish some nomenclature, first note that in a conflict there aretwo sides: the attacker and the defender. The attacker is thetransaction requesting access to a shared memory location. Inpessimistic conflict detection, the attacker is the transaction issuingthe load or load exclusive. In optimistic, the attacker is thetransaction attempting to validate. The defender in both cases is thetransaction receiving the attacker's request.

An Aggressive CM Policy immediately and always retries either theattacker or the defender. In LO, Aggressive means that the attackeralways wins, and so Aggressive is sometimes called committer wins. Sucha policy was used for the earliest LO systems. In the case of EP,Aggressive can be either defender wins or attacker wins.

Restarting a conflicting transaction that will immediately experienceanother conflict is bound to waste work—namely interconnect bandwidthrefilling cache misses. A Polite CM Policy employs exponential backoff(but linear could also be used) before restarting conflicts. To preventstarvation, a situation where a process does not have resourcesallocated to it by the scheduler, the exponential backoff greatlyincreases the odds of transaction success after some n retries.

Another approach to conflict resolution is to randomly abort theattacker or defender (a policy called Randomized). Such a policy may becombined with a randomized backoff scheme to avoid unneeded contention.

However, making random choices, when selecting a transaction to abort,can result in aborting transactions that have completed “a lot of work”,which can waste resources. To avoid such waste, the amount of workcompleted on the transaction can be taken into account when determiningwhich transaction to abort. One measure of work could be a transaction'sage. Other methods include Oldest, Bulk TM, Size Matters, Karma, andPolka. Oldest is a simple timestamp method that aborts the youngertransaction in a conflict. Bulk TM uses this scheme. Size Matters islike Oldest but instead of transaction age, the number of read/writtenwords is used as the priority, reverting to Oldest after a fixed numberof aborts. Karma is similar, using the size of the write-set aspriority. Rollback then proceeds after backing off a fixed amount oftime. Aborted transactions keep their priorities after being aborted(hence the name Karma). Polka works like Karma but instead of backingoff a predefined amount of time, it backs off exponentially more eachtime.

Since aborting wastes work, it is logical to argue that stalling anattacker until the defender has finished their transaction would lead tobetter performance Unfortunately, such a simple scheme easily leads todeadlock.

Deadlock avoidance techniques can be used to solve this problem. Greedyuses two rules to avoid deadlock. The first rule is, if a firsttransaction, T1, has lower priority than a second transaction, T0, or ifT1 is waiting for another transaction, then T1 aborts when conflictingwith T0. The second rule is, if T1 has higher priority than T0 and isnot waiting, then T0 waits until T1 commits, aborts, or starts waiting(in which case the first rule is applied). Greedy provides someguarantees about time bounds for executing a set of transactions. One EPdesign (LogTM) uses a CM policy similar to Greedy to achieve stallingwith conservative deadlock avoidance.

Example MESI coherency rules provide for four possible states in which acache line of a multiprocessor cache system may reside, M, E, S, and I,defined as follows:

Modified (M): The cache line is present only in the current cache, andis dirty; it has been modified from the value in main memory. The cacheis required to write the data back to main memory at some time in thefuture, before permitting any other read of the (no longer valid) mainmemory state. The write-back changes the line to the Exclusive state.

Exclusive (E): The cache line is present only in the current cache, butis clean; it matches main memory. It may be changed to the Shared stateat any time, in response to a read request. Alternatively, it may bechanged to the Modified state when writing to it.

Shared (S): Indicates that this cache line may be stored in other cachesof the machine and is “clean”; it matches the main memory. The line maybe discarded (changed to the Invalid state) at any time.

Invalid (I): Indicates that this cache line is invalid (unused).

TM coherency status indicators (R 132, W 138) may be provided for eachcache line, in addition to, or encoded in the MESI coherency bits. An R132 indicator indicates the current transaction has read from the dataof the cache line, and a W 138 indicator indicates the currenttransaction has written to the data of the cache line.

In another aspect of TM design, a system is designed using transactionalstore buffers. U.S. Pat. No. 6,349,361 titled “Methods and Apparatus forReordering and Renaming Memory References in a Multiprocessor ComputerSystem,” filed Mar. 31, 2000 and incorporated by reference herein in itsentirety, teaches a method for reordering and renaming memory referencesin a multiprocessor computer system having at least a first and a secondprocessor. The first processor has a first private cache and a firstbuffer, and the second processor has a second private cache and a secondbuffer. The method includes the operations of, for each of a pluralityof gated store requests received by the first processor to store adatum, exclusively acquiring a cache line that contains the datum by thefirst private cache, and storing the datum in the first buffer. Upon thefirst buffer receiving a load request from the first processor to load aparticular datum, the particular datum is provided to the firstprocessor from among the data stored in the first buffer based on anin-order sequence of load and store operations. Upon the first cachereceiving a load request from the second cache for a given datum, anerror condition is indicated and a current state of at least one of theprocessors is reset to an earlier state when the load request for thegiven datum corresponds to the data stored in the first buffer.

The main implementation components of one such transactional memoryfacility are a transaction-backup register file for holdingpre-transaction GR (general register) content, a cache directory totrack the cache lines accessed during the transaction, a store cache tobuffer stores until the transaction ends, and firmware routines toperform various complex functions. In this section a detailedimplementation is described.

IBM zEnterprise EC12 Enterprise Server Embodiment

The IBM zEnterprise EC12 enterprise server introduces transactionalexecution (TX) in transactional memory, and is described in part in apaper, “Transactional Memory Architecture and Implementation for IBMSystem z” of Proceedings Pages 25-36 presented at MICRO-45, 1-5 Dec.2012, Vancouver, British Columbia, Canada, available from IEEE ComputerSociety Conference Publishing Services (CPS), which is incorporated byreference herein in its entirety.

Table 3 shows an example transaction. Transactions started with TBEGINare not assured to ever successfully complete with TEND, since they canexperience an aborting condition at every attempted execution, e.g., dueto repeating conflicts with other CPUs. This requires that the programsupport a fallback path to perform the same operationnon-transactionally, e.g., by using traditional locking schemes. Thisputs significant burden on the programming and software verificationteams, especially where the fallback path is not automatically generatedby a reliable compiler.

TABLE 3 Example Transaction Code LHI R0,0 *initialize retry count=0 loopTBEGIN *begin transaction JNZ abort *go to abort code if CC1=0 LT R1,lock *load and test the fallback lock JNZ lckbzy *branch if lock busy .. . perform operation . . . TEND *end transaction . . .  . . .  . . .  .. . lckbzy TABORT *abort if lock busy; this *resumes after TBEGIN abortJO fallback *no retry if CC=3 AHI R0, 1 *increment retry count CIJNLR0,6, fallback *give up after 6 attempts PPA R0, TX *random delay basedon retry count . . . potentially wait for lock to become free . . . Jloop *jump back to retry fallback OBTAIN lock *using Compare&Swap . . .perform operation . . . RELEASE lock . . .  . . .  . . .  . . .

The requirement of providing a fallback path for aborted TransactionExecution (TX) transactions can be onerous. Many transactions operatingon shared data structures are expected to be short, touch only a fewdistinct memory locations, and use simple instructions only. For thosetransactions, the IBM zEnterprise EC12 introduces the concept ofconstrained transactions; under normal conditions, the CPU 114 (FIG. 2)assures that constrained transactions eventually end successfully,albeit without giving a strict limit on the number of necessary retries.A constrained transaction starts with a TBEGINC instruction and endswith a regular TEND. Implementing a task as a constrained ornon-constrained transaction typically results in very comparableperformance, but constrained transactions simplify software developmentby removing the need for a fallback path. IBM's Transactional Executionarchitecture is further described in z/Architecture, Principles ofOperation, Tenth Edition, SA22-7832-09 published September 2012 fromIBM, incorporated by reference herein in its entirety.

A constrained transaction starts with the TBEGINC instruction. Atransaction initiated with TBEGINC must follow a list of programmingconstraints; otherwise the program takes a non-filterableconstraint-violation interruption. Exemplary constraints may include,but not be limited to: the transaction can execute a maximum of 32instructions, all instruction text must be within 256 consecutive bytesof memory; the transaction contains only forward-pointing relativebranches (i.e., no loops or subroutine calls); the transaction canaccess a maximum of 4 aligned octowords (an octoword is 32 bytes) ofmemory; and restriction of the instruction-set to exclude complexinstructions like decimal or floating-point operations. The constraintsare chosen such that many common operations like doubly linkedlist-insert/delete operations can be performed, including the verypowerful concept of atomic compare-and-swap targeting up to 4 alignedoctowords. At the same time, the constraints were chosen conservativelysuch that future CPU implementations can assure transaction successwithout needing to adjust the constraints, since that would otherwiselead to software incompatibility.

TBEGINC mostly behaves like XBEGIN in TSX or TBEGIN on IBM's zEC12servers, except that the floating-point register (FPR) control and theprogram interruption filtering fields do not exist and the controls areconsidered to be zero. On a transaction abort, the instruction addressis set back directly to the TBEGINC instead of to the instruction after,reflecting the immediate retry and absence of an abort path forconstrained transactions.

Nested transactions are not allowed within constrained transactions, butif a TBEGINC occurs within a non-constrained transaction it is treatedas opening a new non-constrained nesting level just like TBEGIN would.This can occur, e.g., if a non-constrained transaction calls asubroutine that uses a constrained transaction internally.

Since interruption filtering is implicitly off, all exceptions during aconstrained transaction lead to an interruption into the operatingsystem (OS). Eventual successful finishing of the transaction relies onthe capability of the OS to page-in the at most 4 pages touched by anyconstrained transaction. The OS must also ensure time-slices long enoughto allow the transaction to complete.

TABLE 4 Transaction Code Example TBEGINC *begin constrained transaction. . . perform operation . . . TEND *end transaction

Table 4 shows the constrained-transactional implementation of the codein Table 3, assuming that the constrained transactions do not interactwith other locking-based code. No lock testing is shown therefore, butcould be added if constrained transactions and lock-based code weremixed.

When failure occurs repeatedly, software emulation is performed usingmillicode as part of system firmware. Advantageously, constrainedtransactions have desirable properties because of the burden removedfrom programmers.

With reference to FIG. 3, the IBM zEnterprise EC12 processor introducedthe transactional execution facility. The processor can decode 3instructions per clock cycle; simple instructions are dispatched assingle micro-ops, and more complex instructions are cracked intomultiple micro-ops. The micro-ops (Uops 232 b) are written into aunified issue queue 216, from where they can be issued out-of-order. Upto two fixed-point, one floating-point, two load/store, and two branchinstructions can execute every cycle. A Global Completion Table (GCT)232 holds every micro-op 232 b and a transaction nesting depth (TND) 232a. The GCT 232 is written in-order at decode time, tracks the executionstatus of each micro-op 232 b, and completes instructions when allmicro-ops 232 b of the oldest instruction group have successfullyexecuted.

The level 1 (L1) data cache 240 is a 96 KB (kilo-byte) 6-way associativecache with 256 byte cache-lines and 4 cycle use latency, coupled to aprivate 1 MB (mega-byte) 8-way associative 2nd-level (L2) data cache 268with 7 cycles use-latency penalty for L1 240 misses. The L1 240 cache isthe cache closest to a processor and Ln cache is a cache at the nthlevel of caching. Both L1 240 and L2 268 caches are store-through. Sixcores on each central processor (CP) chip share a 48 MB 3rd-levelstore-in cache, and six CP chips are connected to an off-chip 384 MB4th-level cache, packaged together on a glass ceramic multi-chip module(MCM). Up to 4 multi-chip modules (MCMs) can be connected to a coherentsymmetric multi-processor (SMP) system with up to 144 cores (not allcores are available to run customer workload).

Coherency is managed with a variant of the MESI protocol. Cache-linescan be owned read-only (shared) or exclusive; the L1 240 and L2 268 arestore-through and thus do not contain dirty lines. The L3 272 and L4caches (not shown) are store-in and track dirty states. Each cache isinclusive of all its connected lower level caches.

Coherency requests are called “cross interrogates” (XI) and are senthierarchically from higher level to lower-level caches, and between theL4s. When one core misses the L1 240 and L2 268 and requests the cacheline from its local L3 272, the L3 272 checks whether it owns the line,and if necessary sends an XI to the currently owning L2 268/L1 240 underthat L3 272 to ensure coherency, before it returns the cache line to therequestor. If the request also misses the L3 272, the L3 272 sends arequest to the L4 (not shown), which enforces coherency by sending XIsto all necessary L3s under that L4, and to the neighboring L4s. Then theL4 responds to the requesting L3 which forwards the response to the L2268/L1 240.

Note that due to the inclusivity rule of the cache hierarchy, sometimescache lines are XI'ed from lower-level caches due to evictions onhigher-level caches caused by associativity overflows from requests toother cache lines. These XIs can be called “LRU XIs”, where LRU standsfor least recently used.

Making reference to yet another type of XI requests, Demote-XIstransition cache-ownership from exclusive into read-only state, andExclusive-XIs transition cache ownership from exclusive into invalidstate. Demote-XIs and Exclusive-XIs need a response back to the XIsender. The target cache can “accept” the XI, or send a “reject”response if it first needs to evict dirty data before accepting the XI.The L1 240/L2 268 caches are store through, but may reject demote-XIsand exclusive XIs if they have stores in their store queues that need tobe sent to L3 before downgrading the exclusive state. A rejected XI willbe repeated by the sender. Read-only-XIs are sent to caches that own theline read-only; no response is needed for such XIs since they cannot berejected. The details of the SMP protocol are similar to those describedfor the IBM z10 by P. Mak, C. Walters, and G. Strait, in “IBM System z10processor cache subsystem microarchitecture”, IBM Journal of Researchand Development, Vol 53:1, 2009, which is incorporated by referenceherein in its entirety.

Transactional Instruction Execution

FIG. 3 depicts example components of an example transactional executionenvironment, including a CPU and caches/components with which itinteracts (such as those depicted in FIGS. 1 and 2). The instructiondecode unit 208 (IDU) keeps track of the current transaction nestingdepth 212 (TND). When the IDU 208 receives a TBEGIN instruction, thenesting depth 212 is incremented, and conversely decremented on TENDinstructions. The nesting depth 212 is written into the GCT 232 forevery dispatched instruction. When a TBEGIN or TEND is decoded on aspeculative path that later gets flushed, the IDU's 208 nesting depth212 is refreshed from the youngest GCT 232 entry that is not flushed.The transactional state is also written into the issue queue 216 forconsumption by the execution units, mostly by the Load/Store Unit (LSU)280, which also has an effective address calculator 236 is included inthe LSU 280. The TBEGIN instruction may specify a transaction diagnosticblock (TDB) for recording status information, should the transactionabort before reaching a TEND instruction.

Similar to the nesting depth, the IDU 208/GCT 232 collaboratively trackthe access register/floating-point register (AR/FPR) modification masksthrough the transaction nest; the IDU 208 can place an abort requestinto the GCT 232 when an AR/FPR-modifying instruction is decoded and themodification mask blocks that. When the instruction becomesnext-to-complete, completion is blocked and the transaction aborts.Other restricted instructions are handled similarly, including TBEGIN ifdecoded while in a constrained transaction, or exceeding the maximumnesting depth.

An outermost TBEGIN is cracked into multiple micro-ops depending on theGR-Save-Mask; each micro-op 232 b (including, for example uop 0, uop 1,and uop2) will be executed by one of the two fixed point units (FXUs)220 to save a pair of GRs 228 into a special transaction-backup registerfile 224, that is used to later restore the GR 228 content in case of atransaction abort. Also the TBEGIN spawns micro-ops 232 b to perform anaccessibility test for the TDB if one is specified; the address is savedin a special purpose register for later usage in the abort case. At thedecoding of an outermost TBEGIN, the instruction address and theinstruction text of the TBEGIN are also saved in special purposeregisters for a potential abort processing later on.

TEND and NTSTG are single micro-op 232 b instructions; NTSTG(non-transactional store) is handled like a normal store except that itis marked as non-transactional in the issue queue 216 so that the LSU280 can treat it appropriately. TEND is a no-op at execution time, theending of the transaction is performed when TEND completes.

As mentioned, instructions that are within a transaction are marked assuch in the issue queue 216, but otherwise execute mostly unchanged; theLSU 280 performs isolation tracking as described in the next section.

Since decoding is in-order, and since the IDU 208 keeps track of thecurrent transactional state and writes it into the issue queue 216 alongwith every instruction from the transaction, execution of TBEGIN, TEND,and instructions before, within, and after the transaction can beperformed out-of order. It is even possible (though unlikely) that TENDis executed first, then the entire transaction, and lastly the TBEGINexecutes. Program order is restored through the GCT 232 at completiontime. The length of transactions is not limited by the size of the GCT232, since general purpose registers (GRs) 228 can be restored from thebackup register file 224.

During execution, the program event recording (PER) events are filteredbased on the Event Suppression Control, and a PER TEND event is detectedif enabled. Similarly, while in transactional mode, a pseudo-randomgenerator may be causing the random aborts as enabled by the TransactionDiagnostics Control.

Tracking for Transactional Isolation

The Load/Store Unit 280 tracks cache lines that were accessed duringtransactional execution, and triggers an abort if an XI from another CPU(or an LRU-XI) conflicts with the footprint. If the conflicting XI is anexclusive or demote XI, the LSU 280 rejects the XI back to the L3 272 inthe hope of finishing the transaction before the L3 272 repeats the XI.This “stiff-arming” is very efficient in highly contended transactions.In order to prevent hangs when two CPUs stiff-arm each other, aXI-reject counter is implemented, which triggers a transaction abortwhen a threshold is met.

The L1 cache directory 240 is traditionally implemented with staticrandom access memories (SRAMs). For the transactional memoryimplementation, the valid bits 244 (64 rows×6 ways) of the directoryhave been moved into normal logic latches, and are supplemented with twomore bits per cache line: the TX-read 248 and TX-dirty 252 bits.

The TX-read 248 bits are reset when a new outermost TBEGIN is decoded(which is interlocked against a prior still pending transaction). TheTX-read 248 bit is set at execution time by every load instruction thatis marked “transactional” in the issue queue. Note that this can lead toover-marking if speculative loads are executed, for example on amispredicted branch path. The alternative of setting the TX-read 248 bitat load completion time was too expensive for silicon area, sincemultiple loads can complete at the same time, requiring many read-portson the load-queue.

Stores execute the same way as in non-transactional mode, but atransaction mark is placed in the store queue (STQ) 260 entry of thestore instruction. At write-back time, when the data from the STQ 260 iswritten into the L1 240, the TX-dirty bit 252 in the L1-directory 256 isset for the written cache line. Store write-back into the L1 240 occursonly after the store instruction has completed, and at most one store iswritten back per cycle. Before completion and write-back, loads canaccess the data from the STQ 260 by means of store-forwarding; afterwrite-back, the CPU 114 (FIG. 2) can access the speculatively updateddata in the L1 240. If the transaction ends successfully, the TX-dirtybits 252 of all cache-lines are cleared, and also the TX-marks of notyet written stores are cleared in the STQ 260, effectively turning thepending stores into normal stores.

On a transaction abort, all pending transactional stores are invalidatedfrom the STQ 260, even those already completed. All cache lines thatwere modified by the transaction in the L1 240, that is, have theTX-dirty bit 252 on, have their valid bits turned off, effectivelyremoving them from the L1 240 cache instantaneously.

The architecture requires that before completing a new instruction, theisolation of the transaction read- and write-set is maintained. Thisisolation is ensured by stalling instruction completion at appropriatetimes when XIs are pending; speculative out-of order execution isallowed, optimistically assuming that the pending XIs are to differentaddresses and not actually cause a transaction conflict. This designfits very naturally with the XI-vs-completion interlocks that areimplemented on prior systems to ensure the strong memory ordering thatthe architecture requires.

When the L1 240 receives an XI, L1 240 accesses the directory to checkvalidity of the XI'ed address in the L1 240, and if the TX-read bit 248is active on the XI'ed line and the XI is not rejected, the LSU 280triggers an abort. When a cache line with active TX-read bit 248 isLRU'ed from the L1 240, a special LRU-extension vector remembers foreach of the 64 rows of the L1 240 that a TX-read line existed on thatrow. Since no precise address tracking exists for the LRU extensions,any non-rejected XI that hits a valid extension row the LSU 280 triggersan abort. Providing the LRU-extension effectively increases the readfootprint capability from the L1-size to the L2-size and associativity,provided no conflicts with other CPUs 114 (FIGS. 1 and 2) against thenon-precise LRU-extension tracking causes aborts.

The store footprint is limited by the store cache size (the store cacheis discussed in more detail below) and thus implicitly by the L2 268size and associativity. No LRU-extension action needs to be performedwhen a TX-dirty 252 cache line is LRU'ed from the L1 240.

Store Cache

In prior systems, since the L1 240 and L2 268 are store-through caches,every store instruction causes an L3 272 store access; with now 6 coresper L3 272 and further improved performance of each core, the store ratefor the L3 272 (and to a lesser extent for the L2 268) becomesproblematic for certain workloads. In order to avoid store queuingdelays, a gathering store cache 264 had to be added, that combinesstores to neighboring addresses before sending them to the L3 272.

For transactional memory performance, it is acceptable to invalidateevery TX-dirty 252 cache line from the L1 240 on transaction aborts,because the L2 268 cache is very close (7 cycles L1 240 miss penalty) tobring back the clean lines. However, it would be unacceptable forperformance (and silicon area for tracking) to have transactional storeswrite the L2 268 before the transaction ends and then invalidate alldirty L2 268 cache lines on abort (or even worse on the shared L3 272).

The two problems of store bandwidth and transactional memory storehandling can both be addressed with the gathering store cache 264. Thecache 264 is a circular queue of 64 entries, each entry holding 128bytes of data with byte-precise valid bits. In non-transactionaloperation, when a store is received from the LSU 280, the store cache264 checks whether an entry exists for the same address, and if sogathers the new store into the existing entry. If no entry exists, a newentry is written into the queue, and if the number of free entries fallsunder a threshold, the oldest entries are written back to the L2 268 andL3 272 caches.

When a new outermost transaction begins, all existing entries in thestore cache are marked closed so that no new stores can be gathered intothem, and eviction of those entries to L2 268 and L3 272 is started.From that point on, the transactional stores coming out of the LSU 280STQ 260 allocate new entries, or gather into existing transactionalentries. The write-back of those stores into L2 268 and L3 272 isblocked, until the transaction ends successfully; at that pointsubsequent (post-transaction) stores can continue to gather intoexisting entries, until the next transaction closes those entries again.

The store cache 264 is queried on every exclusive or demote XI, andcauses an XI reject if the XI compares to any active entry. If the coreis not completing further instructions while continuously rejecting XIs,the transaction is aborted at a certain threshold to avoid hangs.

The LSU 280 requests a transaction abort when the store cache 264overflows. The LSU 280 detects this condition when it tries to send anew store that cannot merge into an existing entry, and the entire storecache 264 is filled with stores from the current transaction. The storecache 264 is managed as a subset of the L2 268: while transactionallydirty lines can be evicted from the L1 240, they have to stay residentin the L2 268 throughout the transaction. The maximum store footprint isthus limited to the store cache size of 64×128 bytes, and it is alsolimited by the associativity of the L2 268. Since the L2 268 is 8-wayassociative and has 512 rows, it is typically large enough to not causetransaction aborts.

If a transaction aborts, the store cache 264 is notified and all entriesholding transactional data are invalidated. The store cache 264 also hasa mark per doubleword (8 bytes) whether the entry was written by a NTSTGinstruction—those doublewords stay valid across transaction aborts.

Millicode-Implemented Functions

Traditionally, IBM mainframe server processors contain a layer offirmware called millicode which performs complex functions like certainCISC instruction executions, interruption handling, systemsynchronization, and RAS. Millicode includes machine dependentinstructions as well as instructions of the instruction set architecture(ISA) that are fetched and executed from memory similarly toinstructions of application programs and the operating system (OS).Firmware resides in a restricted area of main memory that customerprograms cannot access. When hardware detects a situation that needs toinvoke millicode, the instruction fetching unit 204 switches into“millicode mode” and starts fetching at the appropriate location in themillicode memory area. Millicode may be fetched and executed in the sameway as instructions of the instruction set architecture (ISA), and mayinclude ISA instructions.

For transactional memory, millicode is involved in various complexsituations. Every transaction abort invokes a dedicated millicodesub-routine to perform the necessary abort operations. Thetransaction-abort millicode starts by reading special-purpose registers(SPRs) holding the hardware internal abort reason, potential exceptionreasons, and the aborted instruction address, which millicode then usesto store a TDB if one is specified. The TBEGIN instruction text isloaded from an SPR to obtain the GR-save-mask, which is needed formillicode to know which GRs 238 to restore.

The CPU 114 (FIG. 2) supports a special millicode-only instruction toread out the backup-GRs 224 and copy them into the main GRs 228. TheTBEGIN instruction address is also loaded from an SPR to set the newinstruction address in the PSW to continue execution after the TBEGINonce the millicode abort sub-routine finishes. That PSW may later besaved as program-old PSW in case the abort is caused by a non-filteredprogram interruption.

The TABORT instruction may be millicode implemented; when the IDU 208decodes TABORT, it instructs the instruction fetch unit to branch intoTABORT's millicode, from which millicode branches into the common abortsub-routine.

The Extract Transaction Nesting Depth (ETND) instruction may also bemillicoded, since it is not performance critical; millicode loads thecurrent nesting depth out of a special hardware register and places itinto a GR 228. The PPA instruction is millicoded; it performs theoptimal delay based on the current abort count provided by software asan operand to PPA, and also based on other hardware internal state.

For constrained transactions, millicode may keep track of the number ofaborts. The counter is reset to 0 on successful TEND completion, or ifan interruption into the OS occurs (since it is not known if or when theOS will return to the program). Depending on the current abort count,millicode can invoke certain mechanisms to improve the chance of successfor the subsequent transaction retry. The mechanisms involve, forexample, successively increasing random delays between retries, andreducing the amount of speculative execution to avoid encounteringaborts caused by speculative accesses to data that the transaction isnot actually using. As a last resort, millicode can broadcast to otherCPUs 114 (FIG. 2) to stop all conflicting work, retry the localtransaction, before releasing the other CPUs 114 to continue normalprocessing. Multiple CPUs 114 must be coordinated to not causedeadlocks, so some serialization between millicode instances ondifferent CPUs 114 is required.

As previously described, transactional memory is designed to allowdevelopers to mark the portions of their programs that modify the shareddata as being “atomic.” Each atomic block is executed within atransaction: either the whole block executes, or none of it does. Withinthe atomic block, the program can read the shared value without lockingit, perform all the computations it needs to perform, and then write thevalue back. At the end, it commits the transaction. The transactionalmemory system checks to see if the shared data has been modified sincethe atomic operation was started. If it hasn't, the commit may make theupdate and the thread may carry on with its work. If the shared valuehas changed, the transaction is aborted, and the work the threadperformed is rolled back.

One current method of accelerating execution of programs is speculativeexecution. Speculative execution is an optimization technique where acomputer system performs a task that it anticipates may be needed in thefuture. The main idea of speculative execution is to have a processor dowork before it is known whether that work will actually be needed atall. Doing the work ahead of time may prevent a delay that may have tobe incurred later (i.e., by doing the work after it is known that thework is in fact needed). If it turns out the work was not needed afterall, any changes made by the work are reverted and the results areignored. The target is to provide more concurrency if extra resourcesare available. This technique is employed in a variety of areas,including branch prediction in pipelined processors, prefetching memoryand files, and optimistic concurrency control in a database system.

Speculation may lead to accelerated processing by improving cycles perinstruction (CPI). However, within transactions, misspeculated paths maycause the transaction's footprint to increase by introducing additionalread and write accesses to a transaction's read and write sets sincetransactional execution keeps track of all the memory addresses that thetransaction has referenced so that another transaction cannot referencethe same memory location. As such, these unnecessary additions to atransaction's read and/or write sets may lead to either a capacityinduced fail or a false interference with a request from anotherprocessor/transaction. Therefore, it is desirable to eliminate and/orminimize the number of false entries.

In accordance with embodiment, memory operations (i.e., memoryinstructions having operands in memory in a processor pipeline) may onlybe executed when they are in an access-for completion (AFC) mode (i.e.,speculation throttling mode is enabled) and the memory instructions arenext to complete (NTC), or close to complete. Conversely, memoryoperations (i.e., memory instructions having operands in memory in aprocessor pipeline) may be executed speculatively, when they are not inan access-for completion (AFC) (i.e., speculation throttling mode is notenabled). This may eliminate execution of speculative memory accesses.In one aspect of the present embodiment, the requirement for memoryinstructions to be next or close to completion is enforced adaptively,responsive to a speculation-throttle indication. Speculation throttleindication may be responsive to (i.e., including but not limited) to oneor more of the following: an indication in the system that a transactionhas failed and needs to be executed non-speculatively (e.g., forconstrained transactions, or because a transaction is being retried); anexcessive transaction abort rate in the system, or under control of theapplication (e.g. with a perform processor assist (PPA) instruction—orother speculation control instruction—either explicitly selectingspeculation throttling; or in conjunction with a non-specific PPArequest), either after a transaction has failed for the first time, orproactively prior to executing a transaction (e.g., because thetransaction is known to a programmer as triggering excessivespeculation). In different embodiments, a PPA instruction—or other suchinstruction—may permanently enable speculation throttle indication,indicate it for the next transaction only, or select one or the otherunder programmer control.

For illustrative purposes, FIG. 4 depicts a conventional processor 400(i.e., a pipelined processor) with predictor update logic. The processor400 includes, among other things, prediction hardware, registers,caches, decoders, group formation 145 a, an instruction sequencing unit(ISU) 150, a load store unit (LSU) 170, and instruction execution units.In particular, the prediction hardware includes Local Branch HistoryTable (BHT) 110 a, Global Branch History Table (BHT) 110 b, and GlobalSelector 110 c. The prediction hardware is accessed through anInstruction Fetch Address Register (IFAR) 120, which has the address forthe next instruction fetch. In one embodiment, an instruction cache 125fetches a plurality of instructions referred to as a “fetch group”.

The cache and prediction hardware are accessed at approximately the sametime with the same address. If the prediction hardware has predictioninformation available for an instruction in the fetch group, thatprediction is forwarded to the ISU 150, which, in turn, issuesinstructions to units for execution. The prediction may be used toupdate the IFAR 120 in conjunction with branch target calculation andbranch target prediction hardware (such as a link register predictionstack and a count register cache). If no prediction information isavailable, but the instruction decoders find a branch instruction in thefetch group, a prediction is created for that fetch group, stored in theprediction hardware and forwarded to the ISU 150.

The Branch Execution Unit (BRU) 140 a operates in response toinstructions issued to it by the ISU 150. The BRU 140 a has read accessto the condition register file 160. The Branch Execution Unit 140 afurther has access to information stored by the branch scan logic in theBranch Information Queue (BIQ) 142 a, to determine the success of abranch prediction, and is operatively coupled to the instruction fetchaddress register(s) (IFAR) 120 corresponding to the one or more threadssupported by the microprocessor. In accordance with at least oneembodiment, BIQ entries are associated with, and identified by anidentifier, e.g., by a branch tag BTAG. When a branch associated with aBIQ entry is completed, it is so marked. BIQ entries are maintained in aqueue, and the oldest queue entry (entries) is/are de-allocatedsequentially when they are marked as containing information associatedto a completed branch. The BRU 140 a is further operatively coupled tocause a predictor update when the BRU 140 a discovers a branchmisprediction.

When the instruction is executed, the BRU 140 a detects if theprediction is wrong. If so, the prediction needs to be updated. For thispurpose, the processor in FIG. 4 also includes predictor update logic130 a. The predictor update logic 130 a is responsive to an updateindication from Branch Execution Unit 140 a and configured to updatearray entries in one or more of the Local BHT 110 a, Global BHT 110 b,and Global Selector 110 c. The predictor hardware 110 a, 110 b, and 110c may have write ports distinct from the read ports used by theinstruction fetch and prediction operation, or a single read/write portmay be shared. The predictor update logic 130 a may further beoperatively coupled to the link stack 115 a and counter register stack115 b.

Referring now to the condition register file (CRF) 160, the CRF 160 isread-accessible by the BRU 140 a and can be written by execution unitsincluding but not limited to the Fixed Point Unit (FXU) 165 a, FloatingPoint Unit (FPU) 175 and Vector Multimedia eXtension Unit (VMXU) 180.The Condition Register Logic Execution unit (CRL execution) 155 (alsoreferred to as the CRU) and SPR handling logic have read and writeaccess to the Condition Register File (CRF) 160 (access to CRF 160 fromSPR handling logic not shown in the interest of simplifyingillustration). The CRU 155 performs logical operations on the conditionregisters stored in the CRF file 160. The FXU 165 a is able to performwrite updates to the CRF 160.

Referring now to FIG. 5, an exemplary instruction sequencing unit (ISU)150 in accordance with the present embodiment is depicted. Theinstruction sequencing unit (ISU) 150 is the unit of the processor 400(FIG. 4) that is responsible for deciding when to start executing aspecific instruction. The ISU 150 works by receiving one or moreinstructions from an instruction decode unit (IDU) 300 e. The ISU 150may typically capture the received instructions into one or moreinstruction latches 305. Then the ISU 150 may place the instructionsinto one or more issue queues 300 a-300 d. The issue queues 330 a-330 dare lists of pending instructions that need to be executed. Each issuequeue 330 a-330 d has scheduling logic associated with it that maydecide which instruction should be executed next based on respectingdependencies (e.g. making sure instructions only get executed when alltheir inputs are available and with a goal of executing all the oldestinstructions first so they may complete).

There may be a number of issue queues 300 a-300 d as shown with respectto FIG. 5. According to at least embodiment, each type of instruction,such as a fixed point 300 a, a floating point 300 b, a load/store 300 c,or a branch 300 d may have a separate issue queue 300 a-300 d asdepicted in FIG. 5. However according another implementation, multipleissue queues 300 a-d may be present for the same class. For example,there may be two separate floating point issue queues 300 b. In anotherimplementation, multiple units may share one issue queue 300 a-300 d.Once an instruction is selected, it is sent to its correspondingexecution unit 300 f-300 i. The execution units may be a fixed pointunit (FXU) 300 f; a floating point unit (FPU) 300 g; a load store unit(LSU) 300 h; and a branch unit (BRU) 300 i. A global completion table(GCT) 300 j is also depicted. The GCT 300 j is a general trackingmechanism that may track which instructions in the issue queue 300 a-300d have completed. GCT 300 j may receive completion reports 300 k fromthe various execution units 300 f-300 i and may ultimately beresponsible for reallocating resources in a processor 400 (FIG. 4) whenall the oldest instructions have executed in the order of the originalprogram.

Referring now to FIG. 6, an exemplary issue queue of an instructionsequencing unit (ISU) 600, in accordance with the present embodiment isdepicted. According to one implementation, an issue queue 600 may bedesigned to include a list of instructions 502 that are in the issuequeue 600 in conjunction with target register specifiers (RT) 504 andsource register specifiers (RS1) 506 and (RS2) 508. Each of the sourceregister specifiers 506, 508 may be coupled with tag match logic 510,514 and a ready indication 512, 516 as to whether an input operand isready. This may be a way of determining whether an instruction is readyand has all its inputs available (i.e., whether all the inputs have beencomputed for that instruction) and whether that instruction may beselected (i.e., selected instruction 522) by the select logic 518 toexecute out of order.

Furthermore, every time an instruction 502 finishes execution and writesits result to the specified target register RT, the instruction'sresults tag RT corresponding to its result register(s) are broadcast 520to all other instructions. For example, an instruction 502 may write toregister 2 as its target (RT) 508 and another instruction 502 may bewaiting for register 2 as its second input operand (RS2) 508, as such,the select logic 518 may be able to determine that the input is nowavailable and the pending instruction 502 may have become ready toissue. As such, the select logic 518 may be tracking all thedependencies. Additionally, the select logic 518 may look at all theindications in the instruction 502 and, at least in one aspect of thepresent embodiment, may select one or more of the oldest instructionsafter all of the operands are ready to be executed. Then the selectlogic 518 may select those instructions 522 and send the instructions502 to the various execution units 300 f-300 i (FIG. 5).

Currently, when a speculatively executed memory instruction is undone(rolled back) and removed (i.e., flushed) from the processor 400 (FIG.4), the memory instruction may leave behind the memory instruction'saddress in the read and write sets. However, according to embodiment,out of order execution may be throttled (i.e., limited) for memoryinstructions to prevent the excessive inclusion of addresses intransaction read or write sets due to misspeculation when later thespeculation is rolled back (i.e., the speculation is undone). As such,over-speculation may not produce the extraneous memory addresses to readand write sets due to misspeculation. For example, the out of orderexecution throttling may be implemented as a mode in a processor 400(FIG. 4) that may always be enabled or may only be enabled whenexcessive misspeculation is detected. For example, performance may bedegrading by introducing false interferences (i.e., when a collisionbetween two transactions is detected because they refer to the samememory references). As such a mode of adaptiveness may be used inconjunction with at least on embodiment to trigger the new operatingmode of out of order execution throttling or not depending on whetherexcessive interference based on misspeculation may be a problem. Oneadvantage of implementing the out of order execution throttling modeadaptively is to be able to enable and disable the out of orderexecution throttling mode as necessary. For example, limitingspeculation may negatively impact performance, however if there is toomuch speculation, then the excessive speculation may negatively impactperformance as well.

Typically today, processors 400 (FIG. 4) perform speculation toaccelerate the processing by determining which instructions will beexecuted next and then trying to execute them before it may beestablished or proven that those instructions may in fact need to beexecuted. Usually speculation may lead to performance gain because workis started earlier and therefore the work may be completed earlier.However, when excessive speculation is performed, it may interfere withthe progress of other work since later, it may be determined that thespeculatively executed work should not have been performed. As such,there may be no gain in having performed the speculation, but rather anegative impact resulting from the interference with the actual currentwork that is being performed may have occurred. In particular, and inreference to the maintenance of read and write sets associated withtransactions, speculation can lead to the over-indication of atransaction's read and write sets. As such, it may be helpful to detectand limit the speculation in situations such as excessive speculation.One implementation of an adaptive mode may be to speculate aggressivelyinitially, but if too much interference is detected, then the out oforder execution throttling mode (i.e., the speculation throttling mode)may be initiated. However, according to another implementation, thespeculation throttling mode may always be present.

The instruction sequencing unit (ISU) 150 (FIG. 5) may have instructions502 (FIG. 6) coming in and typically those instructions in instructionlatches 305 (FIG. 5) may immediately be dispersed into the issue queues300 a-300 d (FIG. 5) and marked in the GCT 300 j (FIG. 5) so they becomeeligible to be executed out of order. However, some instructions 305(FIG. 5) may need to be executed in order and typically thoseinstructions 305 (FIG. 5) may be marked to execute as next to complete(NTC). As such, when those instructions of instruction latch(es) 305(FIG. 5) marked as NTC arrive in the instruction sequencing unit (ISU)150 (FIG. 6), they may not be dispersed into the issue queues 300 a-300d (FIG. 5) from the instruction latches 305 (FIG. 5), but rather theinstructions may be stored in the instruction latches 305 of theinstruction sequencing unit (ISU) 150 (FIG. 5) and the processor 400(FIG. 4) may wait until all the previous instructions 502 (FIG. 6) inissue queues 300 a-d have been executed. The ISU 150 (FIG. 5) then mayknow that the pending instruction at the top of the ISU 150 ininstruction latch(es) 305 (FIG. 5) is the NTC instruction 502 (FIG. 6)because there are no instructions 502 (FIG. 6) left in any of the issuequeues 300 a-300 d (FIG. 5) and in the GCT 300 j (FIG. 5) whereinstruction tracking of instructions 502 in issues queues 300 a-d isperformed in conjunction with the tracking of the instructions 502 (FIG.6). However, a disadvantage of this method of ensuring in-orderexecution of instructions marked NTC may be that all the instructionsthat are behind an instruction 502 (FIG. 6) that is marked NTC may beprevented from coming into the ISU 150 (FIG. 5). Therefore, when a NTCinstruction is in the instruction latch 305 (FIG. 5), all subsequentinstructions in the program are blocked from executing because theinstruction of instruction latch(es) 305 (FIG. 5) that is marked as NTCacts as a barrier for all future instructions that may be received fromIDU 300 e (and more specifically shown as decode logic 145 a in FIG. 4).As a result, speculation may be limited further by preventing allinstructions following an instruction marked NTC to be executed out oforder with respect to an NTC instruction.

According to at least one embodiment, memory load and store instructionsmay only be executed when the instructions are marked NTC or close tocompletion (e.g., close to NTC). Memory instructions 502 (FIG. 6) withina transaction may be marked as being NTC and may be held at theinstruction latch 305 (FIG. 5) until the GCT 300 j (FIG. 5) is notholding any pending instructions (i.e., the speculation throttling modeis active). As such, dependent instructions 502 (FIG. 6) and furtherinstructions 502 (FIG. 6) behind a NTC or close to NTC instruction 502(FIG. 6) marked instruction 502 (FIG. 6) are queued; therefore,preventing out of order execution (i.e., speculative execution) withrespect to other instructions 502 (FIG. 6) in the program sequence.

However, performance may negatively be impacted since every time amemory load or store instruction is encountered within a transactionwhile the throttling mode is active, all further instructions may beprevented from executing (i.e., those instructions that are behind thatmemory load or store instruction marked as NTC). As a result, the memoryload or store of the transaction that is marked as NTC may stay in theinstruction latch 305 (FIG. 5), and any future instructions may beprevented from coming into the ISU 150 (FIG. 5). As such, otherimplementations of the present embodiment may allow more speculation tooccur while still limiting speculation with respect to memory load andstore instructions.

Referring now to FIG. 7, an operational flowchart 700 illustrating theoperations carried out by a program to improve memory performance whenengaging next to complete (NTC) speculation throttling is depicted. Assuch, more speculation may be allowed within the processor. As describedabove, throttling may be performed by forcing memory instructions to beNTC. In accordance with at least one embodiment, speculation throttlingapplies to memory load and store instructions when they are within atransaction, but not to prefetch or other such memory hint instructionsthat are not tracked as part of transactional read and write sets. Inother embodiments, all memory instructions within a transaction areaffected. NTC speculation throttling may be engaged either by hardware,or under the control of a speculation mode control instruction. Aspeculation mode control instruction may be any assist instruction tohelp transaction execution in accordance with known prior methods or anew instruction. According to at least embodiment, the speculationcontrol instruction may be a transaction speculation control enableinstruction.

At 602, non-memory instructions may be executed, optionally speculatingacross memory operations ahead of time to accelerate execution. Forexample, there may be long running instructions, such as divide,floating point operations, address computation leading to further memoryoperations, or other instructions might benefit from being started earlyif they are independent of the memory instruction. In one embodiment, aspreviously described, ISU 150 (FIG. 4) and instruction latch 305 (FIG.4) may hold an instruction marked as NTC (i.e., an AFC mode settinginstruction that causes an instruction defined value to be set as an AFCmode value.) In one such embodiment, a memory load or store that ismarked as NTC in a transaction when the speculation throttling mode isenabled does not allow speculation across that memory instruction.

In another embodiment, non-memory instructions may be speculated far inadvance, but the memory instructions may not be speculated far inadvance and are allowed to execute only when they are next to complete,or close to complete (in accordance with a definition of close tocomplete for a specific implementation, such as including, but notlimited to: not further than n instructions from completion, not furtherthan n branches from completion, not further than n unresolved branchesfrom completion, for a variety of possible values of n, and including0). For example, all of the non-memory instructions may be executedahead of the memory instructions. Then the memory instructions may beexecuted at the in-order point (i.e., or close to the in-order point toallow for a small amount of speculation).

Therefore, at 604, a memory instruction may be encountered. Then,according to one implementation, at 606, the method may wait for all theprevious non-memory instructions ahead of the memory instruction tocomplete. Once all the previous non-memory instructions have completed,the method may continue at 608 with the execution of that particularmemory instruction that it has encountered at the in-order point (orclose to the in-order point). Then the method may continue back to 602to select other instructions and continue with speculative execution.

In one embodiment previously described with respect to FIG. 6, whenspeculation throttling mode is enabled, memory instructions (and in atleast one embodiment, preferably memory load and store instructions) aremarked as requiring NTC completion. In another embodiment implementingcompute instructions with memory operands, all instructions having atleast one memory read operand are treated as memory load instructions.For example, any operand that has read-accesses memory as an operand(e.g., ADD R2, 2(r4)—i.e., add the content of memory location at r4+2 tothe contents of register r2) should be handled as a load instruction.This includes executed as NTC, but it also includes being executed asclose to completion, after preceding branches have been verified togenerate a prefetching operation for such a memory operand.

In said another embodiment implementing compute instructions with memoryoperands, all instructions having at least one memory write operand aretreated as memory store instructions. In one embodiment operating inconjunction with the ISU structure of FIG. 5 and issue queues 300 a-300d in accordance with the issue queue architecture of FIG. 6, memoryinstruction marked for NTC completion are held in instruction latch 305(FIG. 5) until they are eligible to execute when they are next tocomplete. In such an embodiment, these instructions thereby act as abarrier and prevent new instructions from entering into the ISU 150(FIG. 5). Since there is an instruction already in the instruction latch305 (FIG. 5), the future instructions may be blocked from entering andall future instructions may be prevented from being speculativelyexecuted, leading to significant—and possibly excessive—reduction ofspeculation near memory instructions in transactions, when speculationthrottling mode is enabled.

At least embodiment is designed to enable execution of non-memoryinstructions subsequent to a memory instruction in a transaction subjectto speculation throttling out of order with respect to said memoryinstruction, According to at least one such embodiment, a load and storeissue queue 300 a-300 d (FIG. 5) is modified to issue memoryinstructions only when they are NTC within a transaction when thespeculation throttling mode has been enabled. As such, memoryinstructions in a transaction when the throttling mode is enabled, areno longer held in the instruction latch 305 (FIG. 5), but are allowed toenter the issue queue 300 c (FIG. 5). Furthermore, any instructionsfollowing that memory instruction may be allowed to go to theirrespective issue queues 300 a-300 d (FIG. 5) and execute out of order.To ensure that the memory load and stores may still be executed whenthey are NTC, the load and store issue queue (LSU) may be modified toonly execute transactional loads and stores in-order (when they are attheir in-order point) according to one implementation.

Referring now to FIG. 8, a modified load and store unit (LSU) issue 800queue is depicted. In accordance with one embodiment, the issue queue ofFIG. 8 enables speculating across memory instructions in accordance withthe present embodiment by implementing next-to-complete orclose-to-complete function in the load/store issue queue while allowinginstructions subsequent to so-marked instructions to be entered intotheir respective issue queues and become eligible for execution out oforder with respect to their program order and with respect to memoryinstructions required to execute in order, as previously described withrespect to the optional aspects of step 602 (FIG. 7). Thus, with respectto the modified LSU issue queue depicted in FIG. 8, memory instructionsmarked to execute in-order when they are next-to-complete are notrequired to be held in the instruction latch 305 (FIG. 5).Advantageously, these instructions will not be acting as a barrierpreventing subsequent instructions from entering the ISU by way ofinstruction latch(es) 305, from being dispersed into issue queues 300 a,300 b, 300 d, etc and from becoming eligible for out of order execution.The first modification to the LSU issue queue (as depicted in FIG. 8) istargeting the authority to speculate other instructions which may bedone by allowing memory loads and stores from transactions, where thethrottling is active, to go to their respective issue queue such thatthe instruction latch 305 (FIG. 5) becomes available again to receiveadditional instruction. As such, the issue queue may be modified toenforce that a memory load and store in a transaction may only executedwhen they are non-speculative (i.e., in order and NTC).

In accordance with one implementation of an LSU issue queue enforcingin-order execution for memory load and store instructions subject tospeculation throttling, the LSU issue queue includes two additionalfields associated with each of the instructions in the issue queue. Onefield may be a one bit indication that an instruction is inside atransaction and subject to the speculation throttling mode (inTX) 704.The inTX 704 field may differentiate between an instruction inside atransaction (i.e., inTX 704), from an instruction that is not inside atransaction. The inTX 704 field may also determine when the speculationthrottling mode is not enabled and as such, the instructions may beperformed out of order. Therefore, the inTX 704 may indicate for eachinstruction 502 in the issue queue whether said instruction is in atransaction with throttling mode enabled or not. In one embodiment,indicator inTX 704 is used as an input in the select logic so that anyinstruction that isn't in a transaction with speculation throttlingenabled may always execute out of order in accordance with methodscurrently used.

In at least one embodiment, each GCT entry is uniquely identified by aGlobal Completion Identifier (GCT ID). Furthermore, in an exemplaryembodiment, each issue queue entry further contains a Global CompletionIdentifier (GCT ID) 706 in addition to the inTX 704 field. The GCT ID706 may point to an entry in the global completion table GCT 300 j (FIG.5). The GCT 300 j (FIG. 5) may keep track of all the instructions inprogress and which are the oldest instructions that are about tocomplete successfully. Therefore, an instruction is non-speculative whenit is the oldest instruction in the GCT 300 j (FIG. 5). Although theissue queue may execute instructions as early as possible, the GCT 300 j(FIG. 5) ensures that even if an instruction executes out of order, eachinstruction completes according to the original program order. Aninstruction is at its in-order point (i.e., NTC) when that instructioncorresponds to an instruction (or instruction group) that is at the headof GCT 300 j (FIG. 5) as next to complete.

In accordance with the issue queue design 800 of FIG. 8, an instructionis eligible to be selected as selected instruction 522 by selectionlogic 518 when all its input operands are marked as ready (columns 512and 516) and an instruction is not marked by indicator 704 as being in atransaction with speculation throttling. Alternatively, a transactioncan be selected when indicator 704 is set, and the instruction is theoldest element in issue queue 800 (i.e., at the bottom row), and its GCTID matches the GCT ID of the next instruction (or, instruction group) tocomplete from the GCT. In another embodiment, memory instructions areeligible when they are within the bottom n instructions of issue queue800, and match either the GCT ID at the head that is next to complete,or, optionally, one of a plurality of GCT IDs representing instructions(instruction groups) that are next to complete, or near to complete.

Therefore, when a memory load and store instruction is in a transaction,speculation throttling may be enabled to ensure the instruction isexecuted in order when it is next-to-complete, i.e., when the GCT ID 706the instruction has matches the ID of the next to complete instructionat the head of the GCT 300 j (FIG. 5). As such, that instruction is thenext instruction to complete. The issue queue may receive the next tocomplete GCT ID 722 regarding the next to complete instruction from theGCT table 300 j (FIG. 5) and if that id matches the id 706 of thecurrent instruction then the issue queue may determine that the currentinstruction is the NTC instruction since that instruction is the next inline according to the GCT 300 j (FIG. 5). Summarily, logic to determinein-transaction status, speculation throttling and/or GCT ID matchrepresents “close to complete select enablement logic” 720 whichqualifies candidate instructions 502 from an exemplary load and storeissue queue 300 c to be selected or not selected by select logic 518. Inat least one embodiment, close to complete select enablement logic 720is operatively coupled to a GCT window control register 726 controllingwhen a memory load or store operations is eligible to be issued as beingnext to or close to complete in conjunction with methods of logic 722.

Referring now to FIGS. 9-11, operational flowcharts 800 illustrating thelogic for logic 720. Logic 722 operates in conjunction with a windowcontrol register 726, in accordance with the present embodiment aredepicted. The GCT window control register 726 (FIG. 8) may be set by anapplication or system programmer A special purpose register may be used.For example, using the mtspr instruction of the Power ISA architecturemay be used, e.g., including a transfer to a next to complete window(NTCW) special purpose register from a general purpose register:

-   -   mtspr NTCW, R2

In another embodiment, a load control register instruction may be used.In another embodiment, a suitable value for control register 726 may bedetermined by a manufacturer and may be loaded by a scan chain, fromserial EEPROM, or by other means during system initialization. Inanother embodiment, the window specification is determined at designtime and tangibly represented in circuitry.

Referring now to FIG. 9, a flowchart for logic 720 (FIG. 8) isillustrated in accordance with at least one implementation of thepresent embodiment. At 728, it is determined whether the instruction isin-transaction 704 (FIG. 8), or in another embodiment is in atransaction that is subject to speculation throttling. And if it isfalse (i.e., In−TX=0), then, at 734, the instruction is determined to bea candidate for selection 518 (FIG. 8) as a selected instruction 522(FIG. 8). However, if at 728, it is determined that the instruction isin-transaction 704 (FIG. 8) (i.e., In−TX=1), or in another embodiment isin a transaction that is subject to speculation throttling. then at 730,it is determined whether the instruction GCT ID 706 (FIG. 8) is equal tothe GCT ID from GCT ID 722 (FIG. 8) and if it is true, then, at 736, theinstruction is determined to be a candidate for selection 518 (FIG. 8)as a selected instruction 522 (FIG. 8). If, at 730, it is determinedthat the instruction GCT ID 706 (FIG. 8) is not equal to the GCT ID fromGCT ID 722 (FIG. 8), then the instruction is not a candidate forselection 518 (FIG. 8) as a selected instruction 522 (FIG. 8) at 732.

Referring now to FIG. 10, a flowchart for logic 720 (FIG. 8) isillustrated in accordance with at least one implementation of thepresent embodiment. At 738, it is determined whether the instruction isin-transaction 704 (FIG. 8), or in another embodiment is in atransaction that is subject to speculation throttling. And if it isfalse (i.e., In−TX=0), then at 746, the instruction is determined to bea candidate for selection 518 (FIG. 8) as a selected instruction 522(FIG. 8). However, if at 738, it is determined that the instruction isin-transaction 704 (FIG. 8) (i.e., In−TX=1), or in another embodiment isin a transaction that is subject to speculation throttling. then at 740,it is determined whether the instruction GCT ID 706 (FIG. 8) is greaterthan or equal to the GCT ID from GCT ID 722 (FIG. 8) and if it is false,then, at 748, the instruction is determined not to be a candidate forselection 518 (FIG. 8) as a selected instruction 522 (FIG. 8). However,if at 740, it is determined that the instruction GCT ID 706 (FIG. 8) isgreater than or equal to the GCT ID from GCT ID 722 (FIG. 8), then at742, it is determined whether the instruction GCT ID 706 (FIG. 8) isless than or equal to the sum of the GCT ID from GCT ID 722 (FIG. 8) thevalue of GCT window 726 (FIG. 8), i.e., whether it is in a window of ninstructions in a first embodiment where a GCT ID uniquely identifiesinstruction, or whether it is in a window of n instruction groups in aanother embodiment where a GCT ID uniquely identifies instructiongroups, where n represents the value corresponding to the window sizestored in GCT window 726, and if it is false, then, at 750, theinstruction is determined not to be a candidate for selection 518 (FIG.8) as a selected instruction 522 (FIG. 8). However, if at 742, it isdetermined that the instruction GCT ID 706 (FIG. 8) is less than orequal to the GCT ID from GCT ID 722 (FIG. 8)+GCT window 726 (FIG. 8),then at 744 the instruction is a candidate for selection 518 (FIG. 8) asa selected instruction 522 (FIG. 8).

Referring now to FIG. 11, a flowchart for logic 720 (FIG. 8) withpending instruction window is illustrated in accordance with at leastone implementation of the present embodiment. At 752, it is determinedwhether the instruction is in-transaction 704 (FIG. 8), or in anotherembodiment is in a transaction that is subject to speculationthrottling. And if it is false (i.e., In−TX=0), then, at 758, theinstruction is determined to be a candidate for selection 518 (FIG. 8)as a selected instruction 522 (FIG. 8). However, if at 752, it isdetermined that the instruction is in-transaction 704 (FIG. 8) (i.e.,In−TX=1), or in another embodiment is in a transaction that is subjectto speculation throttling. Then at 754, it is determined whether theinstruction is within oldest n instructions 502 (FIG. 8)—where n isselected by window control register 726 (FIG. 8) and if it is true thenat 760, the instruction is determined to be a candidate for selection518 (FIG. 8) as selected instruction 522 (FIG. 8). If at 754, theinstruction is determined not to be within oldest n instructions 502(FIG. 8)—where n is selected by window control register 726 (FIG. 8),then at 756, the instruction is determined not to be a candidate forselection 518 (FIG. 8) as selected instruction 522 (FIG. 8).

Referring now to FIG. 12, a modified load and store unit (LSU) issuequeue 900 optimized for allowing other instructions to speculate acrossmemory instructions, in accordance with the present embodiment isdepicted. According to such an alternate embodiment, a processor may bebuilt that forces all memory instructions to be executed in order. Assuch, the tracking of the dependency in accordance with columns 510-516(FIG. 8) can be omitted because the next to complete (oldest)instruction in the global completion table is necessarily ready toissue, because there are no instructions prior to such an instructionfor whose results the next to complete instruction might wait. Sincethere are no prior instructions, this instruction cannot be waiting onany prior instruction. However, such an embodiment that always choosesthe oldest memory load and store instruction may not be adapted to everperform out of order memory instructions. Therefore, if the ability ofallowing out of order instructions outside the transaction is desirable,e.g., when speculation throttling is not enabled, or when a memoryinstruction is outside of a transaction, then this embodiment does notprovide the desired function. However, this embodiment may be more costeffective as compared to other possible implementations and, therefore,may be more desirable to implement if the ability of allowing out oforder instructions outside the transactions is not a requirement.

In the previously explained embodiments, execution of memory load andstore instructions in a transaction have been limited to executing onlywhen the instructions are NTC (i.e., no other instruction before thememory load and store instruction inside the transaction) whilethrottling speculation is active. However, limiting the execution ofinstructions to only when the instruction is NTC may impose asignificant degradation on performance since speculation is generallythought to be an advantage to execution.

As such, one embodiment may reduce the performance penalty by notpreventing all speculation, but by preventing the types of speculationthat are most prone to leading to excessive speculation and as such, endup adding memory addresses to a transaction's read and write setsunnecessarily. For example, branch misprediction is one of the mostfrequent misspeculations that excessively adds unnecessary memorylocations to the transactions' read and write sets.

Therefore, according to one embodiment with speculation throttlingactive, memory load and stores may be held and not executed until thebranch prediction has been verified as correct for all branchespreceding a memory load or store instruction that (i.e., all the priorbranches have successfully executed without a branch misprediction). Abranch, or a branch prediction is verified (or validated, or resolved)when a branch execution executes, finds no misprediction, and completes.This is different than the previously described method of preventingmemory load and store instructions from executing until the instructionsare the NTC (i.e., they are the oldest instruction, or within a fewinstructions of NTC) and there are no other instructions waiting toexecute before that instruction executes in that memory instructions canbe executed significantly out of order as long as branch predictionscorresponding to any branches preceding such memory instructions havebeen verified.

According to one implementation, memory instruction load and stores maybe allowed to execute as soon as any branches preceding that memory loador store instruction have been verified as having correctly beenpredicted. This may prevent the negative impact of misspeculation when abranch misprediction occurs while still providing some flexibility tostart memory loads and stores out of order relative to some otherinstructions.

In accordance with embodiment, each memory load or store instructioninside a transaction when throttling speculation is enabled isassociated with a preceding branch tag (BTAG) 902 (FIG. 10). In thisembodiment, each branch, according to at least one embodiment, may beidentified by a tag or pointer that uniquely identifies the branch. Theidentifier may be used to store, for each memory load or storeoperation, the associated branch that comes before that memoryoperation. In one embodiment, issue logic is extended such that a memoryload or store (that has been marked to be within a transaction whenspeculation throttling has been enabled) is only executed when thebranch that comes before that memory load or store has been validated(i.e., that branch prediction has been verified). Each instruction'sBTAG field 902 (FIG. 10) is further associated with a ready field 906(FIG. 10) responsive to BTAG tag match 904 (FIG. 10) logic which is afurther input to instruction selection logic 522. In accordance withthis embodiment, the memory instructions BTAG field 902 (FIG. 10) isoperatively coupled with tag match logic 904 (FIG. 10) that matches amemory instructions' BTAG field (corresponding to its immediatelypreceding branch instruction's BTAG) to the BTAG IDs of completedbranches. In one preferred embodiment, a memory instruction 502 (FIG.10) is marked as ready 906 (FIG. 10) only when the branch instructionidentified by the BTAG 902 (FIG. 10) and all preceding branches havecompleted.

In accordance with an aspect of the present embodiment, executing memoryload or store instructions only after preceding branch instructions havecompleted is beneficial because, as previously described, when amisprediction occurs, the system may flush (i.e., eliminate) themisspeculated load or store instruction. However, according to currentmethods, flushing the misspeculated load or store instruction from theprocessor does not update the transaction's memory read and write set.As such the associated entries in the transaction's memory read andwrite set are not removed. Therefore a correction of a misspeculationinside a processor does not correct the memory read or write setsmaintained for a transaction. Conversely, when it is determined thatthere is no misspeculation possible on all branches preceding a memoryload or store instruction since they have been correctly predicted, thenit may be safe to start executing a memory load or store instructioneven if there are other instructions before it. There may be otherevents that may conceptually lead to a misspeculation, such as afloating point exception or an external exception; however, such amisspeculation would abort the entire transaction. Therefore, such anexception has no effect on the memory load or store instructions and thetransaction's read or write sets, but rather typically leads to thecomplete termination of the transaction because restarting executioninside a transaction would not be feasible.

As described above, each memory load or store instruction is associatedwith the branch that precedes that load or store instruction. As such,when it is determined that the identified and preferably all itspreceding branch instructions have been verified (i.e., the branchinstruction has been successfully executed), then the current memoryload or store instruction may be executed.

According to one embodiment, a queue of pending branch executions (i.e.,the branch information queue (BIQ) 142 a (FIG. 4). The BIQ 142 a (FIG.4) may track the branches in the processor that assigns and keeps a tagor reference number for each of the branches in flight and as branchescomplete, the oldest branch at the head of the BIQ is removed from thequeue once it has completed. Branches may complete out of order,however, they are taken from the queue in-order and the oldest branch isremoved from the queue first. When a branch completes and is taken outof the BIQ 142 a (FIG. 4), the tag is the number of the branch that isthe oldest, in-order branch that just completed. Additionally, all thebranches prior to that branch completed as well since that branch is theoldest branch to complete. As such, the oldest tag number completing inthe BIQ 142 a (FIG. 4) may be broadcast by BTAG broadcast logic 908(FIG. 10) to mark as ready all memory load and store instructionassociated with the currently retiring BTAG (provided other dependences,such as the operand dependences 512 (FIG. 5) and 516 (FIG. 5) are met aswell). Additionally, when the oldest entry out of the BIQ 142 a (FIG. 4)is broadcast, then all the branches prior to that branch have also beenvalidated. Thus, when broadcasting the retiring BIQ entry's BTAG, abroadcast BTAG implies not only that the branch identified by thebroadcast BTAG has completed, but rather that all the branches up tothat point have been validated for correctness of prediction.

Referring now to FIG. 13, a modified load and store unit (LSU) issuequeue 1000 speculating across memory instructions, in accordance withthe present embodiment is depicted. As previously described, there maybe instructions 502; registers 506, 508 used by those instructions; tagmatch logic 510, 514 that ensures that each of the input registers to aninstruction has been computed before that instruction can be selected.Additionally, all available register tags may be broadcast 702.Associated with each of these memory instructions 502 is a branch tag(BTAG) 902. The branch corresponding to BTAG 902 should be completedbefore the current instruction 502 may be allowed to execute. As such,there is a BTAG field 902 and a tag match logic 904 that compares theBTAG field 902 against the broadcast BTAG 908. In one embodiment, thebroadcast BTAG corresponds to the BTAG ID of the oldest reclaimed BIQentry 908. Therefore, when a branch instruction corresponding to BTAG908 completes, all the memory instructions that immediately follow thatbranch instruction become eligible to execute (i.e., ready 906) sincethe source of misspeculation has been cleared (i.e., prevented) andmisspeculation may no longer be a concern.

Another alternate embodiment may be implemented with the use of matrixschedulers since issue queues based on wake-up and select logic arecomplex and may be expensive to implement. Current methods ofimplementing matrix schedulers have been incorporated by referenceherein: Masahiro Goshima et al., A High-Speed Dynamic InstructionScheduling Scheme for Superscalar Processors, Proceedings of the 34thInternational Symposium on Microarchitecture MICRO 2001; Mary D. Brownet al., Select-Free Instruction Scheduling Logic, Proceedings of the34th International Symposium on Microarchitecture MICRO 2001, andSamantika Subramaniam and Gabriel H. Loh, Store Vectors for ScalableMemory Dependence Prediction and Scheduling, Conference onHigh-Performance Computer Architecture HPCA 12.

Referring now to FIG. 14, a matrix scheduler 1100 is depicted. In anembodiment, a matrix scheduler represents and tracks dependencies as amatrix of instructions. As such, all the rows 1004 are instructions inthe issue queue and the columns 1002 are pending instructions. For eachinstruction 1006, the matrix depicts which instructions are in flight(i.e., the instructions in the rows 1004) and which instructions thatinstruction may depend upon (i.e., the instructions in the columns1002). For example, when instruction “n+4” depends on instruction J, a“1” may be depicted and when “n+4” depends on another instruction (i.e.,J+4), then another “1” may be depicted. Therefore, a single bit may beat the intersection of the instructions that are waiting to be issued1004 and the instructions that are in the processor 1002 to determine ifthere is a dependency. In accordance with one embodiment of a matrixscheduler, when an instruction J issues, it clears dependence bits ofall instructions depending on that instruction. When all dependence bitsin a row associated with an instruction have cleared, that instruction,e.g., “n+3” becomes eligible to issue. As such, this implementation maybe faster and easier to implement than the wake-up based tag-matchingissue queue embodiment previously described.

Now referring to FIG. 15, there is depicted a matrix scheduler 1200executing instructions only after all prior instructions have executed,in accordance with the present embodiment. In accordance with such anembodiment, dependence information in a matrix scheduler is created toimplement the speculation throttling policy responsive to memoryinstructions (preferably load and store instructions) in a transactionwhen speculation throttling is enabled. In accordance with logicdirected at creating dependence information to implement a speculationthrottling policy when memory instructions are executed only when theyare next to complete, dependence information logic marks candidatememory instructions as dependent on all their preceding instructions.FIG. 12 depicts the operation of a matrix scheduler in conjunction withthis policy with an in-transaction memory instruction 1102 that issubject to speculation throttling and the instruction is marked 1104 tobe dependent on all preceding instructions. As such, the markedinstruction 1104 may only be executed after all the other instructionshave been executed that are marked with a “1” in the columns of the rowcorresponding to instructions 1106. Therefore the policies that havepreviously described may be easier to implement in accordance with amatrix scheduler such as the one depicted in FIG. 12. For example, if anin transaction memory load or store needed to be forced to execute onlyafter all the other instructions before that instruction 1104 areexecuted, then a “1” may be marked in the columns 1106 of all thepreceding instructions. Therefore, the instruction in the row, that hasbeen marked 1104, may only be allowed to execute after all its previousinstructions 1106 have been issued and have cleared their associatedcolumn. As such, an NTC policy that executes a memory instruction onlyafter all the previous instructions have executed, may be implemented bymarking the row 1104 as being dependent on all preceding instructions.

Now referring to FIG. 16, there is depicted a matrix scheduler 1300executing instructions only after all prior branch instructions haveexecuted, in accordance with the present embodiment. There may be anin-transaction memory instruction 1202 subject to speculation throttlingand an instruction may be marked 1204 to be dependent on all real datadependencies (shown as “1” in columns 1216) and on all precedingbranches (shown as “1” in columns 1206). As such, the marked instruction1204 may only be executed after all the other instructions have beenexecuted that are marked with a “1” and are represented in the columnsof the dependence matrix of the matrix scheduler. The figure shows anexemplary in-transaction memory operation to be dependent on two branchinstructions in columns 1206. Therefore the policies that havepreviously described may be easier to implement in accordance with amatrix scheduler such as the one depicted in FIG. 13. For example, if anin transaction memory load or store 1202 needed to be forced to executeonly after all the branches prior to that instruction have beenvalidated, then a “1” may be marked in the columns 1206 corresponding toall the preceding branches. Therefore, the instruction in the row, thathas been marked 1204, may only be allowed to execute after all thebranches prior 1206 to that instruction 1204 have been validated intheir associated column.

In accordance with one embodiment, these policies are implemented aspart of the dependence analysis logic. Other than determining which bitsto set in accordance with the policies described herein, there may be noadditional modification needed to implement a matrix scheduler tosupport the executions previously described with respect to FIGS. 15-16.

Speculation throttling, limiting the execution of memory loads or storeswithin a transaction, may not always need to be implemented immediately.Therefore, speculation throttling, limiting the execution of memoryloads or stores within a transaction, may be controlled by an adaptivespeculation throttling mode which may correspond to misspeculations.

Referring now to FIG. 17, an adaptivity flowchart 1400 illustrating theoperations carried out by a processor to determine whether thespeculation throttling mode is enabled in accordance with the presentembodiment is depicted. The speculation throttling mode, which wasexplained in detail with respect to FIG. 7 above, may be enabled by aprogram with statically inserted instructions to enable and disable thespeculation throttling mode. For example, by a compiler or programmer,based on heuristics, program analysis, program instrumentation, orpreviously observed program behavior. In other embodiments, speculationthrottling mode may also be enabled autonomously by hardware, or acombination of hardware and software.

At 1302, it may be determined whether the speculation throttling modewas enabled. If it is determined at 1302 that the mode is enabled, thenin accordance with the flowchart described with respect to FIG. 7, allnon-memory instructions may be executed, optionally speculating acrossmemory instructions at 602. When a memory instructed is encountered at604, all previous instructions may be completed at 606 and then at 608,the memory instruction may be performed. If it is determined at 1302,that the speculation throttling mode is not enabled, then at 1304, allmemory and non-memory instructions may be executed speculatively.

Referring now to FIG. 18, a flowchart 1500 illustrating the operationsto enable speculation throttling mode based on dynamic behavior inaccordance with the present embodiment is depicted. The speculationthrottling mode may be enabled by a program dynamically based on dynamicfailure counts (i.e., misspeculations), or enabled by a library,middleware, an operating system, a hypervisor, a firmware, a dynamicoptimization component, a millicode, or other responsive to dynamicfailure patterns or counts in a system. Additionally, the speculationthrottling mode may be enabled by hardware responsive to failurepatterns or counts without requiring the need of programmer intervention(i.e., the processor may enable the speculation throttling mode).

At 1402, it is determined whether excessive interference has beenencountered, For example, whether excessive misspeculations haveoccurred. If it is determined at 1402 that excessive interference hasbeen encountered, then at 1404, throttling speculation may be enabled.Conversely, if at 1402 it is determined that excessive interference hasnot been encountered, then at 1406, throttling speculation may bedisabled.

Referring now to FIG. 19, a flowchart 1600 illustrating the operationsfor a perform processor assist (PPA) instruction in accordance with thepresent embodiment is depicted. At 1502, it may be determined whetherthe speculation throttling mode is indicated to be enabled. If, at 1502,it is determined that the speculation throttling mode is not enabled andnot needed, then at 1504, current PPA methods may be implemented.

If at 1502, it is determined the speculation throttling mode is to beconfigured to be enabled or disabled, then at 1506, the speculationthrottling mode may either be enabled (i.e., if it is not enabledalready) or disabled (i.e., if it is enabled and it needs to bedisabled).

Referring now to FIG. 20, an exemplary flowchart 1700 where throttlinginstruction execution in a transaction operating in a processorconfigured to execute memory instructions out-of-order in a pipelinedprocessor, wherein memory instructions are instructions for accessingoperands in memory 1702 is depicted. Included is executing, by theprocessor, instructions of a transaction 1704 including determiningwhether the transaction is in throttling mode 1706 and based on thetransaction being in throttling mode, executing memory instructionsin-program-order 1708. Based on the transaction not-being in throttlingmode, executing memory instructions out-of-program order 1710. Alsoincluded is executing memory instructions in throttling mode based onthe memory instructions being in predefined near-to-complete position inthe processor pipeline 1712 and executing memory instructions inthrottling mode based on the memory instructions being not-speculativerelative to resolution of older branch instructions 1714. Also includedis executing a throttling mode setting instruction, the executingcausing an instruction defined value to be set as a throttling modevalue to indicate throttling mode 1716. The memory instructions being inthe predefined near-to-complete position consist of at least one ofmemory instructions that are not further than a predefined number ofinstructions from completion, memory instructions that are not furtherthan a predefined number of branches from completion, and memoryinstructions that are not further than a predefined number of unresolvedbranches from completion 1718. The memory instructions being in thepredefined near-to-complete position consist of memory instructions thatare next-to-complete 1720. The throttling mode is engaged by a hardwareor by a program 1722. Thee hardware includes at least one of a library,a middleware, an operating system, a hypervisor, a firmware, a dynamicoptimization component, a millicode, or a hardware responsive to dynamicfailure pattern or counts in a system 1724.

FIG. 21 is a block diagram 1800 of internal and external components ofthe hardware and the software of the computer environment according tothe present embodiment. It should be appreciated that FIG. 21 providesonly an illustration of one implementation and does not imply anylimitations with regard to the environments in which differentembodiments may be implemented. Many modifications to the depictedenvironments may be made based on design and implementationrequirements.

Data processing system 800, 900 is representative of any electronicdevice capable of executing machine-readable program instructions. Dataprocessing system 800, 900 may be representative of a smart phone, acomputer system, PDA, or other electronic devices. Examples of computingsystems, environments, and/or configurations that may represented bydata processing system 800, 900 include, but are not limited to,personal computer systems, server computer systems, thin clients, thickclients, hand-held or laptop devices, multiprocessor systems,microprocessor-based systems, network PCs, minicomputer systems, anddistributed cloud computing environments that include any of the abovesystems or devices.

A computer and network server may include respective sets of internalcomponents 800 and external components 900. Each of the sets of internalcomponents 800 includes one or more processors 820, one or morecomputer-readable RAMs 822 and one or more computer-readable ROMs 824 onone or more buses 826, and one or more operating systems 828 and one ormore computer-readable tangible storage devices 830. The one or moreoperating systems 828 and software programs are stored on one or more ofthe respective computer-readable tangible storage devices 830 forexecution by one or more of the respective processors 820 via one ormore of the respective RAMs 822 (which typically include cache memory).In the embodiment illustrated in FIGS. 3-6, each of thecomputer-readable tangible storage devices 830 is a magnetic diskstorage device of an internal hard drive. Alternatively, each of thecomputer-readable tangible storage devices 830 is a semiconductorstorage device such as ROM 824, EPROM, flash memory or any othercomputer-readable tangible storage device that can store a computerprogram and digital information.

Each set of internal components 800 also includes a R/W drive orinterface 832 to read from and write to one or more portablecomputer-readable tangible storage devices 936 such as a CD-ROM, DVD,memory stick, magnetic tape, magnetic disk, optical disk orsemiconductor storage device. A software program can be stored on one ormore of the respective portable computer-readable tangible storagedevices 936, read via the respective R/W drive or interface 832 andloaded into the respective hard drive 830.

Each set of internal components 800 also includes network adapters orinterfaces 836 such as a TCP/IP adapter cards, wireless wi-fl interfacecards, or 3G or 4G wireless interface cards or other wired or wirelesscommunication links A software program in a client computer can bedownloaded to a client computer from an external computer via a network(for example, the Internet, a local area network or other, wide areanetwork) and respective network adapters or interfaces 836. From thenetwork adapters or interfaces 836, the software program in clientcomputer is loaded into the respective hard drive 830. The network maycomprise copper wires, optical fibers, wireless transmission, routers,firewalls, switches, gateway computers and/or edge servers.

Each of the sets of external components 900 can include a computerdisplay monitor 920, a keyboard 930, and a computer mouse 934. Externalcomponents 900 can also include touch screens, virtual keyboards, touchpads, pointing devices, and other human interface devices. Each of thesets of internal components 800 also includes device drivers 840 tointerface to computer display monitor 920, keyboard 930 and computermouse 934. The device drivers 840, R/W drive or interface 832 andnetwork adapter or interface 836 comprise hardware and software (storedin storage device 830 and/or ROM 824).

Aspects of the present embodiment have been described with respect toblock diagrams and/or flowchart illustrations of methods, apparatus(system), and computer program products according to the presentembodiment. It will be understood that each block of the flowchartillustrations and/or block diagrams, and combinations of blocks in theflowchart illustrations and/or block diagrams, can be implemented bycomputer instructions. These computer instructions may be provided to aprocessor of a general purpose computer, special purpose computer, orother programmable data processing apparatus to produce a machine, suchthat instructions, which execute via the processor of the computer orother programmable data processing apparatus, create means forimplementing the functions/acts specified in the flowchart and/or blockdiagram block or blocks.

The aforementioned programs can be written in any combination of one ormore programming languages, including low-level, high-level,object-oriented or non object-oriented languages, such as Java,Smalltalk, C, and C++. The program code may execute entirely on theuser's computer, partly on the user's computer, as a stand-alonesoftware package, partly on the user's computer and partly on a remotecomputer, or entirely on a remote computer or server. In the latterscenario, the remote computer may be connected to the user's computerthrough any type of network, including a local area network (LAN) or awide area network (WAN), or the connection may be made to an externalcomputer (for example, through the Internet using an Internet serviceprovider). Alternatively, the functions of the aforementioned programscan be implemented in whole or in part by computer circuits and otherhardware (not shown).

The foregoing description of various embodiments of the presentembodiment has been presented for purposes of illustration anddescription. It is not intended to be exhaustive or to limit theembodiment to the precise form disclosed. Many modifications andvariations are possible. Such modifications and variations that may beapparent to a person skilled in the art of the embodiment are intendedto be included within the scope of the embodiment as defined by theaccompanying claims.

Various embodiments of the embodiment may be implemented in a dataprocessing system suitable for storing and/or executing program codethat includes at least one processor coupled directly or indirectly tomemory elements through a system bus. The memory elements include, forinstance, local memory employed during actual execution of the programcode, bulk storage, and cache memory which provide temporary storage ofat least some program code in order to reduce the number of times codemust be retrieved from bulk storage during execution.

Input/Output or I/O devices (including, but not limited to, keyboards,displays, pointing devices, DASD, tape, CDs, DVDs, thumb drives andother memory media, etc.) can be coupled to the system either directlyor through intervening I/O controllers. Network adapters may also becoupled to the system to enable the data processing system to becomecoupled to other data processing systems or remote printers or storagedevices through intervening private or public networks. Modems, cablemodems, and Ethernet cards are just a few of the available types ofnetwork adapters.

The present invention may be a system, a method, and/or a computerprogram product. The computer program product may include a computerreadable storage medium (or media) having computer readable programinstructions thereon for causing a processor to carry out aspects of thepresent invention.

The computer readable storage medium can be a tangible device that canretain and store instructions for use by an instruction executiondevice. The computer readable storage medium may be, for example, but isnot limited to, an electronic storage device, a magnetic storage device,an optical storage device, an electromagnetic storage device, asemiconductor storage device, or any suitable combination of theforegoing. A non-exhaustive list of more specific examples of thecomputer readable storage medium includes the following: a portablecomputer diskette, a hard disk, a random access memory (RAM), aread-only memory (ROM), an erasable programmable read-only memory (EPROMor Flash memory), a static random access memory (SRAM), a portablecompact disc read-only memory (CD-ROM), a digital versatile disk (DVD),a memory stick, a floppy disk, a mechanically encoded device such aspunch-cards or raised structures in a groove having instructionsrecorded thereon, and any suitable combination of the foregoing. Acomputer readable storage medium, as used herein, is not to be construedas being transitory signals per se, such as radio waves or other freelypropagating electromagnetic waves, electromagnetic waves propagatingthrough a waveguide or other transmission media (e.g., light pulsespassing through a fiber-optic cable), or electrical signals transmittedthrough a wire.

Computer readable program instructions described herein can bedownloaded to respective computing/processing devices from a computerreadable storage medium or to an external computer or external storagedevice via a network, for example, the Internet, a local area network, awide area network and/or a wireless network. The network may comprisecopper transmission cables, optical transmission fibers, wirelesstransmission, routers, firewalls, switches, gateway computers and/oredge servers. A network adapter card or network interface in eachcomputing/processing device receives computer readable programinstructions from the network and forwards the computer readable programinstructions for storage in a computer readable storage medium withinthe respective computing/processing device.

Computer readable program instructions for carrying out operations ofthe present invention may be assembler instructions,instruction-set-architecture (ISA) instructions, machine instructions,machine dependent instructions, microcode, firmware instructions,state-setting data, or either source code or object code written in anycombination of one or more programming languages, including an objectoriented programming language such as Java, Smalltalk, C++ or the like,and conventional procedural programming languages, such as the “C”programming language or similar programming languages. The computerreadable program instructions may execute entirely on the user'scomputer, partly on the user's computer, as a stand-alone softwarepackage, partly on the user's computer and partly on a remote computeror entirely on the remote computer or server. In the latter scenario,the remote computer may be connected to the user's computer through anytype of network, including a local area network (LAN) or a wide areanetwork (WAN), or the connection may be made to an external computer(for example, through the Internet using an Internet Service Provider).In some embodiments, electronic circuitry including, for example,programmable logic circuitry, field-programmable gate arrays (FPGA), orprogrammable logic arrays (PLA) may execute the computer readableprogram instructions by utilizing state information of the computerreadable program instructions to personalize the electronic circuitry,in order to perform aspects of the present invention.

Aspects of the present invention are described herein with reference toflowchart illustrations and/or block diagrams of methods, apparatus(systems), and computer program products according to embodiments of theinvention. It will be understood that each block of the flowchartillustrations and/or block diagrams, and combinations of blocks in theflowchart illustrations and/or block diagrams, can be implemented bycomputer readable program instructions.

These computer readable program instructions may be provided to aprocessor of a general purpose computer, special purpose computer, orother programmable data processing apparatus to produce a machine, suchthat the instructions, which execute via the processor of the computeror other programmable data processing apparatus, create means forimplementing the functions/acts specified in the flowchart and/or blockdiagram block or blocks. These computer readable program instructionsmay also be stored in a computer readable storage medium that can directa computer, a programmable data processing apparatus, and/or otherdevices to function in a particular manner, such that the computerreadable storage medium having instructions stored therein comprises anarticle of manufacture including instructions which implement aspects ofthe function/act specified in the flowchart and/or block diagram blockor blocks.

The computer readable program instructions may also be loaded onto acomputer, other programmable data processing apparatus, or other deviceto cause a series of operational steps to be performed on the computer,other programmable apparatus or other device to produce a computerimplemented process, such that the instructions which execute on thecomputer, other programmable apparatus, or other device implement thefunctions/acts specified in the flowchart and/or block diagram block orblocks.

The flowchart and block diagrams in the Figures illustrate thearchitecture, functionality, and operation of possible implementationsof systems, methods, and computer program products according to variousembodiments of the present invention. In this regard, each block in theflowchart or block diagrams may represent a module, segment, or portionof instructions, which comprises one or more executable instructions forimplementing the specified logical function(s). In some alternativeimplementations, the functions noted in the block may occur out of theorder noted in the figures. For example, two blocks shown in successionmay, in fact, be executed substantially concurrently, or the blocks maysometimes be executed in the reverse order, depending upon thefunctionality involved. It will also be noted that each block of theblock diagrams and/or flowchart illustration, and combinations of blocksin the block diagrams and/or flowchart illustration, can be implementedby special purpose hardware-based systems that perform the specifiedfunctions or acts or carry out combinations of special purpose hardwareand computer instructions.

Although one or more examples have been provided herein, these are onlyexamples. Many variations are possible without departing from the spiritof the present embodiment. For instance, processing environments otherthan the examples provided herein may include and/or benefit from one ormore aspects of the present embodiment. Further, the environment neednot be based on the z/Architecture®, but instead can be based on otherarchitectures offered by, for instance, IBM®, Intel®, Sun Microsystems,as well as others. Yet further, the environment can include multipleprocessors, be partitioned, and/or be coupled to other systems, asexamples.

As used herein, the term “obtaining” includes, but is not limited to,fetching, receiving, having, providing, being provided, creating,developing, etc.

The flow diagrams depicted herein are just examples. There may be manyvariations to these diagrams or the steps (or operations) describedtherein without departing from the spirit of the embodiment. Forinstance, the steps may be performed in a differing order, or steps maybe added, deleted, or modified. All of these variations are considered apart of the claimed embodiment.

Although preferred embodiments have been depicted and described indetail herein, it will be apparent to those skilled in the relevant artthat various modifications, additions, substitutions and the like can bemade without departing from the spirit of the embodiment, and these are,therefore, considered to be within the scope of the embodiment, asdefined in the following claims.

What is claimed is:
 1. A computer implemented method for throttlinginstruction execution in a transaction operating in a processorconfigured to execute memory instructions out-of-order in a pipelinedprocessor, wherein memory instructions are instructions for accessingoperands in memory, the method comprising: executing, by the processor,instructions of a transaction comprising: executing non-memoryinstructions without engaging speculation throttling, whereinspeculation throttling reduces the number of speculative transactionalexecutions; identifying a first set of memory instructions within thetransaction, wherein the memory instructions must be executed in aparticular order; based on identifying a first set of memoryinstructions within the transaction, engaging speculation throttling;identifying a second set of memory instructions within the transaction,wherein the memory instructions must be executed out-of-program order;and based on identifying a second set of memory instructions within thetransaction, executing the second set of memory instructions withoutengaging speculation throttling.
 2. The method according to claim 1,further comprising: determining whether a position of the memoryinstructions within the processor pipeline is within a thresholddistance of being executed; and executing memory instructions engagingspeculation throttling based on the memory instructions beingnot-speculative relative to resolution of older branch instructions. 3.The method according to claim 1, further comprising: executing athrottling mode setting instruction, the executing causing aninstruction defined value to be set as a throttling mode value toindicate throttling mode.
 4. The method according to claim 2, whereinthe determined memory instructions within the processor pipeline iswithin a threshold distance of being executed consist of at least one ofmemory instructions that are not further than a predefined number ofinstructions from completion, memory instructions that are not furtherthan a predefined number of branches from completion, and memoryinstructions that are not further than a predefined number of unresolvedbranches from completion.
 5. The method according to claim 2, whereinthe determined memory instructions within the processor pipeline iswithin a threshold distance of being executed consist of memoryinstructions that are next in a queue to be completed.
 6. The methodaccording to claim 1, where in the throttling mode is engaged by ahardware or by a program.